Systems and methods for allowing incremental journaling

ABSTRACT

In one embodiment, systems and methods are provided for incremental journaling. In one embodiment, order-independent operations are journaled incrementally for the same storage location. In one embodiment, partially ordered operations are journaled incrementally for the same storage location. In one embodiment, order-independent operations and partially ordered operations are journaled incrementally for the same storage location. In one embodiment, incremental journaling is used to update data that represents accounting, ctime, and parity.

LIMITED COPYRIGHT AUTHORIZATION

A portion of the disclosure of this patent document includes materialwhich is subject to copyright protection. The copyright owner has noobjection to the facsimile reproduction by anyone of the patent documentor the patent disclosure as it appears in the Patent and TrademarkOffice patent file or records, but otherwise reserves all copyrightswhatsoever.

CROSS-REFERENCED APPLICATIONS

This application was filed on the same day as the followingapplications: U.S. application Ser. No. 11/506,597, entitled “SystemsAnd Methods For Providing Nonlinear Journaling,” U.S. application Ser.No. 11/507,073, entitled “Systems And Methods For Providing NonlinearJournaling,” and U.S. application Ser. No. 11/507,070, entitled “SystemsAnd Methods For Providing Nonlinear Journaling,” all of which are herebyincorporated by reference in their entirety herein.

FIELD OF THE INVENTION

This invention relates generally to systems and methods of nonlinearjournaling.

BACKGROUND

The increase in processing power of computer systems has ushered in anew era in which information is accessed on a constant basis. Oneresponse has been to distribute processing requests across multiplenodes or devices. A distributed architecture allows for more flexibleconfigurations with respect to factors such as speed, band-widthmanagement, and other performance and reliability parameters.

A distributed architecture system may provide for numerous storageoperations across many multiple storage devices. In general, recordinginformation regarding these storage operations in a journal—that is,journaling storage operation information—may provide an effective frontend for storage operations, which may support the reliability andcooperation of nodes within a system. Problems may arise with currentsystems for journaling data in both distributed and non-distributedsystems.

Because of the foregoing challenges and limitations, there is an ongoingneed to improve the manner in which computer systems journal storageoperations, such as data writes, on storage devices.

SUMMARY OF THE INVENTION

The systems and methods generally relate to the nonlinear journaling ofdata. In one embodiment, a nonlinear method of journaling data beingwritten to a storage unit is provided. The method may include storing aplurality of groups of data in a journal located in persistent storage;storing information about the location and status of each of saidplurality of groups of data; providing a data structure linking thestored groups of data and the information about each of said groups ofdata; and providing for the unlinking of any data group and itscorresponding stored information, without regards to the order in whichthe data group was stored in the journal.

In another embodiment, a system for journaling data writes into a linkeddata structure in conjunction with storage of the data writes isprovided. The system may include a data storage unit; persistent memoryassociated with said data storage unit; and a program module configuredto journal in said persistent memory data writes to said data storageunit, said data writes comprising data to be written to said datastorage unit and respective locations in said data storage unit to writesaid data; wherein said program module is configured to journal saiddata writes nonlinearly.

In another embodiment, a journal data structure for recording datawrites in conjunction with storage of data writes in a data storage unitis provided. The journal data structure may include a plurality ofjournal blocks comprising data; a plurality of block descriptors, eachblock descriptor comprising a link to at least one of said journalblocks and at least one respective address in the data storage unitassociated with said at least one journal block; and a plurality oftransaction descriptors, each transaction descriptor comprising a linkto at least one of said block descriptors.

In another embodiment, a method of journaling data into a linked datastructure in conjunction with storage of the data in a data storage unitis provided. The method may include journaling data in persistentmemory; journaling in persistent memory a location in the data storageunit, said location corresponding to said data; and linking said dataand said location.

In another embodiment, a method of journaling data into a linked datastructure in conjunction with storage of the data in a data storage unitis provided. The method may include journaling data in persistentmemory; journaling in persistent memory a location in the data storageunit, said location corresponding to said data; and associating saiddata and said location; wherein said data and said location are recordedin nonlinear locations in said persistent memory.

In another embodiment, a method of journaling data for a data storageunit with multiple storage devices is provided. The method may includejournaling data to be written on a plurality of storage devices of adata storage unit; determining when one of said storage devices isunavailable; and keeping the data journaled for said one storage devicewhile said storage device is unavailable for storage.

In another embodiment, a system supporting per-drive journal replay fora data storage unit with multiple storage devices is provided. Thesystem may include a plurality of storage devices; persistent memory;and a program module configured to keep in said persistent memoryjournal data corresponding to a subset of said plurality of storagedevices; wherein said subset of storage devices is temporarilyunavailable to store said journal data.

In another embodiment, a networked cluster of data storage nodescooperating to execute transactions that are global to the networkedcluster of storage nodes is provided. The networked cluster may includea plurality of storage nodes configured to be connected in a network;and a plurality of journal modules, each one of said storage nodeshaving a different one of said plurality of journal modules associatedtherewith, the journal modules configured to record on the storage nodesdata associated with global transactions; wherein the recorded data issufficient to recreate the transactions when necessary.

In another embodiment, a method of journaling data associated withglobal transactions in a distributed data storage system is provided.The method may include journaling data in persistent memory that isassociated with a data storage unit in the distributed data storagesystem, said data associated with a transaction that is global to thedistributed data storage system; wherein journaling said data comprisesrecording information sufficient to recreate the transaction.

In anther embodiment, a networked cluster of data storage nodescooperating to execute transactions that are global to the networkedcluster of data storage nodes is provided. The networked cluster mayinclude a plurality of data storage nodes configured to be connected ina network; a plurality of persistent memory allocations, each one ofsaid data storage nodes having a different one of said plurality ofpersistent memory allocations associated therewith; and at least onejournal program module, the at least one journal program moduleconfigured to record on a subset of the persistent memory allocationsdata associated with transactions that are global to the networkedcluster of data storage nodes; wherein the recorded data is sufficientto recreate the global transactions when necessary.

In another embodiment, a method of journaling data in a storage unit ofa distributed storage system to provide a shadow buffer in the eventthat the distributed system aborts a transaction is provided. The methodmay include journaling first data, said first data associated with afirst transaction that the distributed storage system has committed towrite, said first data designated to be written to a storage location,but said first data has not yet been written to said storage location;journaling second data, said second data associated with a secondtransaction that the distributed storage system has not yet committed towrite, said second data designated to be written to said storagelocation; and preserving said first data for purposes of restoring saidfirst data, in the event that the distributed storage system aborts saidsecond transaction.

In another embodiment, a system that journals data for a data storageunit that provides a shadow buffer in the event that a transactionaborts is included. The system may include a data storage unit; a memorybuffer, said memory buffer associated with a location on said datastorage unit; persistent memory, said persistent memory associated withsaid data storage unit; and a program module configured to journal thefirst data in said persistent memory from said memory buffer, andfurther configured to preserve, after said memory buffer is overwrittenwith second data, the first data in said persistent memory until one ofthe following conditions is met: a transaction associated with thesecond data commits the second data to being stored at the location onsaid data storage unit and, in the event the second data is notcommitted, the first data has been stored to said data storage unit.

In another embodiment, a method of journaling data in a data storageunit of a distributed data storage system to provide both a journalfunction and a shadow buffer function is provided. The method includingkeeping a first data in a memory buffer, the memory buffer associatedwith a location in a data storage unit; journaling the first data inpersistent memory; overwriting the memory buffer with second data beforethe first data is stored in the data storage unit; and preserving thefirst data in the persistent memory until after receiving an indicationthat it may be erased.

In another embodiment, a method conditioning the removal of data from ajournal upon verification that the data has been reliably stored on astorage device of a data storage unit is provided. The method mayinclude journaling data designated to be written to a location on astorage device associated with a data storage unit, said data storagedevice being one of a plurality of storage devices associated with saiddata storage unit; directing the storage device to record the data atthe location; and selectively removing the data from the journal in anorder based upon a communication from the storage device that the datahas been stored at the location without regards to the order in whichthe data was stored.

In another embodiment, a method conditioning the removal of data from ajournal based upon a determination of the least recently used data for aparticular drive is provided. The method may include journaling datadesignated to be written to a location on a storage device associatedwith a data storage unit, said data storage device being one of aplurality of storage devices associated with said data storage unit;directing the storage device to record the data at the location; andselectively removing the data from the journal based upon adetermination that the data is the least recently used datacorresponding to the storage device.

In another embodiment, a system of journaling data that removes datafrom the journal upon synchronizing the contents of a data storage unitis provided. The system may include a data storage unit associated witha plurality of storage devices; persistent memory associated with saiddata storage unit; a synchronization module configured to activelysynchronize contents of said persistent memory with contents of saidplurality of storage devices; and a journaling module configured tojournal data in the persistent memory, and further configured to removedata from said persistent memory after the synchronization modulesynchronizes.

In another embodiment, a system for atomically updating a journal for adata storage unit is provided. The system may include persistent memory;and a program module configured to update a journal located in saidpersistent memory with atomic operations; wherein the journal is not inan inconsistent state following a write failure to update the journal.

In another embodiment, a method of building atomically a journal for adata storage unit is provided. The method may include recording data inpersistent memory, the data being too large to be recorded in a singleatomic operation; and building a journal in the persistent memory, thejournal comprising the data; wherein the journal is built with atomicoperations such that the journal does not comprise only a portion of thedata.

In another embodiment, a concurrent transaction subsystem for a journalas a reliable high-speed front end for disk writes is provided. Theconcurrent transaction subsystem may include a module configured towrite at least one data block to a journal, wherein the journalcomprises an allocation of persistent storage, and wherein the at leastone data block is associated with a location on a memory; wherein themodule is further configured to write at least one delta element to thejournal, wherein the at least one delta element is associated with atleast one data operation that is one of the following: order independentand partially ordered; and wherein the at least one delta element isassociated with the location on the memory.

In another embodiment, a method of implementing a concurrent transactionsubsystem for a journal as a reliable high-speed front end for diskwrites is provided. The method may include writing at least one datablock to a journal, wherein the journal comprises an allocation ofpersistent storage, and wherein the at least one data block isassociated with a location on a memory; and writing at least one deltaelement to the journal, wherein the at least one delta element isassociated with at least one data operation that is one of thefollowing: order independent and partially ordered; and wherein the atleast one delta element is associated with the location on the memory.

For purposes of this summary, certain aspects, advantages, and novelfeatures of the invention are described herein. It is to be understoodthat not necessarily all such advantages may be achieved in accordancewith any particular embodiment of the invention. Thus, for example,those skilled in the art will recognize that the invention may beembodied or carried out in a manner that achieves one advantage or groupof advantages as taught herein without necessarily achieving otheradvantages as may be taught or suggested herein.

BRIEF DESCRIPTION OF THE DRAWINGS

FIGS. 1A, 1B, and 1C illustrate embodiments of a networked cluster ofnodes that journals data.

FIGS. 2A and 2B illustrate embodiments of a networked cluster of nodesthat coordinates storage transactions.

FIGS. 3A, 3B, and 3C illustrate embodiments of data structures forimplementing a nonlinear journal.

FIGS. 4A, 4B, 4C, 4D, and 4E illustrate state diagrams of embodiments ofa nonlinear journal.

FIGS. 5A, 5B, 5C, and 5D illustrate embodiments of building, modifying,and removing nonlinear journal data structures.

FIGS. 6A, 6B, and 6C illustrate flowcharts of embodiments of building,modifying, and removing nonlinear journal data structures.

FIGS. 7A and 7B illustrate flowcharts of one embodiment of rebuildingand replaying, respectively, transactions.

FIG. 8 illustrates embodiments of data structures for implementing anonlinear journal configured to include a shadow buffer.

FIGS. 9A, 9B and 9C illustrate embodiments of keeping a shadow buffer ina nonlinear journal.

FIG. 10 illustrates a flowchart of one embodiment of implementing shadowbuffers in a nonlinear journal.

FIGS. 11A and 11B illustrate embodiments of data structures forimplementing a nonlinear journal capable of handling concurrenttransactions.

FIGS. 12A, 12B-1, 12B-2, 12C-1, 12C-2, 12D-1, and 12D-2 illustrateembodiments of implementing concurrent transactions in a nonlinearjournal.

FIG. 13 illustrates a flowchart of one embodiment of implementingconcurrent transactions in a nonlinear journal.

FIG. 14 illustrates a flowchart of one embodiment of collapsing deltadata structures in a nonlinear journal.

FIG. 15 illustrates one embodiment of combining commit order independentoperations with partially ordered block writes.

These and other features will now be described with reference to thedrawings summarized above. The drawings and the associated descriptionsare provided to illustrate embodiments of the invention and not to limitthe scope of the invention. Throughout the drawings, reference numbersmay be reused to indicate correspondence between referenced elements. Inaddition, the first digit of each reference number generally indicatesthe figure in which the element first appears.

DETAILED DESCRIPTION OF PREFERRED EMBODIMENTS

Systems and methods which represent one embodiment of an exampleapplication of the invention will now be described with reference to thedrawings. Variations to the systems and methods which represent otherembodiments will also be described.

For purposes of illustration, some embodiments will be described in thecontext of a distributed file system. The present invention is notlimited by the type of environment in which the systems and methods areused, however, and systems and methods may be used in otherenvironments, such as, for example, other file systems, otherdistributed systems, the Internet, the World Wide Web, a private networkfor a hospital, a broadcast network for a government agency, and aninternal network for a corporate enterprise, an Intranet, a local areanetwork, a wide area network, a wired network, a wireless network, andso forth. Some of the figures and descriptions, however, relate to anembodiment of the invention wherein the environment is that of adistributed file system. It is also recognized that in otherembodiments, the systems and methods may be implemented as a singlemodule and/or implemented in conjunction with a variety of other modulesand the like. Moreover, the specific implementations described hereinare set forth in order to illustrate, and not to limit, the invention.The scope of the invention is defined by the appended claims.

One example of a distributed file system, in which embodiments ofsystems and methods described herein may be implemented, is described inU.S. patent application Ser. No. 10/007,003 entitled “Systems andMethods for Providing a Distributed File System Utilizing Metadata toTrack Information About Data Stored Throughout the System,” filed Nov.9, 2001 which claims priority to application No. 60/309,803 filed Aug.3, 2001, U.S. patent application Ser. No. 10/281,467 entitled “Systemsand Methods for Providing A Distributed File System Incorporating aVirtual Hot Spare,” filed Oct. 25, 2002, and U.S. patent applicationSer. No. 10/714,326 entitled “Systems And Methods For Restriping FilesIn A Distributed File System,” filed Nov. 14, 2003, which claimspriority to Application No. 60/426,464, filed Nov. 14, 2002, all ofwhich are hereby incorporated by reference herein in their entirety.

I. Overview

In general, embodiments of the invention relate to nonlinear journaling.In computing systems, a journal may provide a reliable high-speedfront-end for disk writes, which implements coherent transactionslocally and rewrites data after power or write failure. Hard-disk drivestypically have an asynchronous write cache. Requests for disk writes areacknowledged immediately by the drive; however, the data is not actuallywritten to stable storage for an unknown amount of time after the writehas been acknowledged. Without a journal, the drive is in an unknownstate after a power or write failure. In other words, the contents ofthe disks (within the drive) are no longer reliable because it is notknown whether the data was actually written to the disks before thefailure occurred. A journal may be used to return the drive back into aknown state. A system equipped with a journal records in the journal thedisk writes to the drive over a period of time. After a power or a writefailure, the system accesses the journal to reissue the writes beforethe drive is used again. Some journal systems simply keep a fixed amountof data to replay in a ring buffer, expiring the oldest—or LeastRecently Used (LRU)—data to make space for new writes. Such animplementation may, however, stall the journal as discussed below.

In addition to providing reliable writes, a journal system may alsoimplement transactions. A collection of blocks may be written under asingle transaction, writing either all of the blocks or no blocks. Thisfeature may also be used in a global transaction system to implementcluster wide transactions. In a cluster wide transaction, journals onmultiple nodes are synchronized such that a transaction on each nodeassociated with a single global transaction either commits or aborts.

It will be appreciated by one skilled in the art that there are manyways to journal data writes. Some problems may arise when implementingjournaling. First, journals may be subject to a single stuck transactionthat halts the entire journal. Some journals are ring-based, in whichthe journal expires the oldest data to make space for new writes. In aring-based system, a single stuck transaction can halt the entirejournal because a transaction that has not been flushed cannot be freedfor a new write, and a new write cannot occur until this stucktransaction completes if it is the oldest data. This may cause journalspace deadlock problems, requiring an extensive journal process forclearing out deadlocks.

Second, some journal systems may not have support for when a node goescompletely offline. As discussed above, many journal systems arering-based, which provide no support for long-lived blocks in a journal.In a ring-based journal, the allocated journal space is systematicallybeing overwritten, which means that no journal block—a data block in thejournal—may last longer in the journal than a full cycle through thejournal. If the journal includes data to be written to a drive that iscurrently down, the journal will likely expire that data prior to thedrive being brought back up. Because bringing back a down drive requiresholding on to blocks for replay over a relatively long period of time,previous journal systems did not support replaying data for down drives.

Third, some journals may require journal escaping. In a linear journalsystem, journal blocks may be written to appear as journal meta blocks.For example, a malicious user could write a journal block that mimicsmetadata in order to corrupt the system. In order to prevent the systemfrom treating the journal block as metadata, linear journal systems mayrequire “journal escaping.” Journal escaping is a procedure whereby thesystem scans data blocks before they are written in the journal, markingin the journal those data blocks that appear to be metadata. Duringsystem replay, marked journal blocks are treated as journal blocks, notmetadata. Escaping journals for proper replay may be a performanceproblem because this may be the only place the data is touched afterbeing given to the journal. Journal escaping can pollute the CPU's cacheand use extra CPU time.

Fourth, some journal systems may require complicated replay. Becauselinear journals overwrite the oldest data with new writes, linearjournals typically have partially written journal blocks. Moreover,linear journals that implement transactions typically have partiallyoverwritten transactions as a result of the linear overwrite.Furthermore, linear journals may suffer from ambiguous head and tailpointers, which identify the beginning of the ring of data for replay.Partially written blocks, partially overwritten transactions, andambiguous head and tail pointers add complexity to using the journal,causing development problems and bugs.

Fifth, some journals implement sloppy disk flushing. Following a powerfailure, the contents of the journal are replayed. Hopefully, thecontents of the journal are sufficient to rewrite anything that was lostin a drive's write cache. This may cause a problem if there is notenough data storage for a particular drive.

Sixth, some previous journal systems implement shadow buffers in systemmemory. A shadow buffer holds a copy of data written to a particularjournal block so that it may be used in the event of an abort of asuccessive write to the same block. If a write is aborted, then thein-memory block corresponding to the journal block of the aborted writeshould be rewritten with the previous value of the buffer before theaborted write. The journal block could be read from disk but may cause alag in performance. Another solution is to keep an in-memory shadowbuffer, but this may also cause a lag in performance

Seventh, some journal systems only support one writer modifying a sameblock at a time. If there are multiple writers on the block, each writerhas to wait for the previous transaction to commit before it couldproceed.

Embodiments of a nonlinear journal (NLJ) described herein may addressone or more of the above issues.

Although the drawings and accompanying description generally describeembodiments of the invention in terms of a distributed architecture,embodiments of the invention are not limited in application to anetworked cluster of computers. Systems and methods for nonlinearjournaling may also be implemented in non-distributed systems.

II. Exemplary Distributed System

FIGS. 1A, 1B, and 1C illustrate embodiments of a networked cluster ofnodes that journals data. FIG. 1A illustrates a high-level diagram ofone embodiment of a distributed system 100 that may journal storagetransactions between nodes 102. In the illustrated embodiment,distributed system 100 includes six nodes 102. Three of these nodes 102are designated as being involved in a particular transaction. Node-1 102and Node-6 102 are Participants of the transaction and are designated asP₂ and P₁, respectively. Node-3 102 is designated as the Initiator ofthe transaction, the Coordinator for the transaction, and a Participantof the transaction. Because Node-3 102 is an initiator and coordinatoras well as a participant, it is a Shared Participant designated asP_(S). The meaning of the terms “Participant,” Shared Participant,”“Initiator” and “Coordinator” in a global transaction system will bediscussed in further detail below with respect to FIGS. 2A and 2B.Although the illustrated embodiment illustrates a single transaction,involving just three nodes 102, it will be appreciated by one skilled inthe art that the distributed system 100 may execute numeroustransactions, which may include various suitable combinations of thenodes 102 as Participants, Shared Participants, Coordinators, andInitiators.

Although in the illustrated embodiment nodes 102 are arranged in a fullyconnected network topology, in other embodiments of the invention, nodes102 may be arranged in other topologies, including, but not limited to,the following topologies: ring, mesh, star, line, tree, bus topologies,and so forth. It will be appreciated by one skilled in the art thatvarious network topologies and/or combinations thereof may be used toimplement different embodiments of the invention. In addition, it isrecognized that nodes 102 may be connected directly, indirectly, or acombination of the two, and that all of the nodes 102 may be connectedusing the same type of connection or one or more different types ofconnections. It is also recognized that in other embodiments, adifferent number of nodes may be included in the cluster, such as, forexample, 2, 16, 83, 6, 883, 10,000, and so forth.

In one embodiment, the nodes 102 are interconnected through abi-directional communication link where messages are received in theorder they are sent. In one embodiment, the link comprises a“keep-alive” mechanism that quickly detects when nodes or other networkcomponents fail, and the nodes are notified when a link goes up or down.In one embodiment, the link includes a Transmission Control Protocol(TCP) connection. In other embodiments, the link includes an SessionDescription Protocol (SDP) connection over Infiniband, a wirelessnetwork, a wired network, a serial connection, Internet Protocol (IP)over FibreChannel, proprietary communication links, connection baseddatagrams or streams, and/or connection based protocols.

In one embodiment of the invention, nodes 102 are individual datastorage units for a distributed file system. FIG. 1B illustrates oneembodiment of a node 102 in a distributed file system, such as thedistributed system 100. In the illustrated embodiment, the nodes 102include hard-disk drives 150, processor 151, system memory 152,persistent memory 156, and system modules 157. In the illustratedembodiment, hard-disk drives 150 store respective portions of thedistributed file system assigned to the respective node 102 indistributed system 100. Processor 151 executes system modules 157 thatmanage the distributed file system, such as read and/or write requests.While executing system modules 157, processor 151 utilizes system memory152 to store data including, for example, data corresponding to read andwrite requests. In the illustrated embodiment, nodes 102 include arespective persistent storage 156, which may be used to store a journalof, for example, data writes to hard-disk drives 150. In the illustratedembodiment, the system modules 157 include journal modules 158, whichmanage the operation of the journal stored in persistent storage 156. Ajournal module 158 may also refer to a journal subsystem, or itscomponents, processes, procedures, and so forth. System modules 157 mayalso include global transaction modules, message processing modules,write modules, read modules, drive sync modules, least recently used(LRU) modules, participant modules, initiator modules, coordinatormodules and so forth.

Although in the illustrated embodiment system modules 157 areillustrated as a separate component, the system modules 157 are programinstructions that may be stored in a variety of suitable locations,including, for example, local partitions on hard-disk drives 150 ordedicated storage devices. Moreover, although the nodes 102 individuallystore their, portions of the distributed file system in an array oftwelve hard-disk drives 150, in other embodiments the nodes 102 mayinclude a different-sized array of hard-disk drives 150, includingpossibly a single hard-disk drive 150. Furthermore, although embodimentsof the invention are generally described with respect to storage devicesbased on hard-disk drives, other embodiments may be implemented onsystems including alternative forms of secondary storage, such as solidstate disks (or drives), random access memory (RAM) disks, Flash disks,combinations of the same, and suitable equivalents. Similarly,embodiments of the invention may include storage devices with variousimplementations of system memory 152, including primary storage based onstatic RAM (SRAM), non-volatile RAM (NVRAM), dynamic RAM (DRAM),combinations of the same, and suitable equivalents. It will beappreciated by one skilled in the art how to implement embodiments ofthe invention on storage systems using suitable alternative devices forprimary and/or secondary storage.

In the illustrated embodiment, a journal of disk writes is stored inpersistent memory 156. Persistent memory, as described herein, may referto memory devices whose contents remain stable despite power failure tothe device. For example, a hard-disk drive, such as one of the hard-diskdrives 150, is an example of persistent storage. Hard-disk drives retaintheir contents, even in the absence of a power supply. Hard-disk drivesdo not, however, have efficient random access. Relatively long seektimes limit the advantageous use of hard-disk drives for journalstorage. Although a hard-disk drive may be used to store a journal, insome embodiments nonvolatile random access memory (NVRAM) is preferred.Flash memory, for example, has faster access times in comparison withhard-disk drives. One disadvantage of flash memory, however, is itsrelatively limited lifecycle. In one embodiment, persistent memory 156is battery-backed RAM. If persistent memory 156 loses power, the backupbattery maintains its persistent state. Battery-backed RAM has theadvantage of efficient access time, long lifecycle, and persistentstate, making it a suitable source of persistent memory 156 for storinga journal. Because battery-backed RAM can lose its memory contents inthe event that the battery fails, persistent memory 156 includes notonly those storage mediums that maintain their contents without anypower; such as a hard-disk drive, but may also include storage mediumswith suitable power-supply backups. Persistent memory 156 may alsoinclude magnetic random access memory (MRAM), which has access time andlifecycle advantages of battery-backed RAM without the need for a backuppower supply. It will be appreciated by one skilled in the art thatpersistent memory 156 may include many suitable forms of nonvolatilememory, including, for example, magnetic random access memory (MRAM),Flash RAM, battery-backed RAM, combinations of the same, and suitableequivalents.

FIG. 1C illustrates one embodiment of a hard-disk drive 150. In theillustrated embodiment, hard-disk drives 150 include a disk controller170 and disk platters 174. The disk controller 170 processes read andwrite requests to the respective disk platter 174, and it includes adrive cache 172 that, for example, buffers data blocks waiting to bewritten to the disk platters 174. Because the drive cache 172 mayinclude volatile storage medium, the contents of the drive cache 172,may be lost without power. A problem arises with disk drives that do notdistinguish between acknowledging data written to the drive cache 172and acknowledging data written to a disk platter 174. If a hard-diskdrive 150 only acknowledges writing data blocks to the drive cache 172,without later acknowledging writing the data blocks to the disk platters174, a power failure may result in an inconsistent state because it isuncertain whether the data block was written to the respective diskplatter 174 before the power failure caused the contents of the drivecache 172 to be lost. That is, following a power failure, a node 102 maynot be able to confirm whether a write request was actually written tothe respective disk platter 174 because the respective hard-disk drive150 acknowledged writing data that was actually lost in the drive cache172 as a result of the power failure. One purpose for journaling data inpersistent memory, such as persistent memory 156, is to preserve data inthe event of a power failure to a hard-disk drive 150, resulting in aninconsistent state due to unacknowledged data loss in the respectivedrive cache 172. That is, one function of a journal in distributedsystem 100 may be to keep data backed up in the event of an inconsistentstate caused by a power failure affecting respective hard-disk drives150.

In embodiments described herein, groups of data in the distributed filesystem are organized into data blocks. Conceptually, a data block may beany size of data, such as a single bit, a byte, a gigabyte, or evenlarger. In general, a data block is the smallest logical unit of datastorage in the file system. In some embodiments, a file system may usedata block sizes that are different from the native block size of adisk. For example, a disk may have a native size of 512 bytes, but afile system may address 4096 bytes or 8192 bytes. In one embodiment, ajournal may handle data blocks in the native size of the file system,not the disk. In another embodiment, a journal may handle data blocksbased on the native size of the disk. In general, the terms “diskblock,” “cache block,” “memory block,” and “journal block,” describedbelow with reference to FIG. 1C, refer to locations in memory devicescorresponding to the size of a “data block.” One skilled in the art willappreciate that file systems may be implemented with many suitable datablock sizes, including, but not limited to, 512 bytes, 4096 bytes, and8192 bytes. In some embodiments, the block size may be configurable. Itwill be further appreciated that, although the illustrated embodimentillustrates a single data block size, file systems may be implementedwith variably sized data blocks.

In the illustrated embodiment, there are several physical storage spacesin a node 102 that correspond to the same data block on a particulardisk platter 174 of a respective hard-disk drive 150. Blocks of data onthe disk platters are referred to as disk blocks 178. Blocks of data ina drive cache 172 correspond to respective disk blocks 178 and arereferred to as cache blocks 176. Although a drive cache 172 may not havea cache block 176 for every disk block 178, a cache block 176 may map tomultiple disk blocks 178, though not at the same time. Typically, a diskblock 178 is written from the copy of the data stored in thecorresponding cache block 176. That is, a cache bock 176 temporarilystores the data value to be written to a corresponding disk block 178.

There may also be blocks of data in the system memory 152 thatcorrespond to a particular disk block 172; these portions of systemmemory 152 are referred to as memory blocks 180. A memory block 180 maystore the present value of a corresponding disk block 178, following,for example, a successful read operation of the respective disk block178. Alternatively, a memory block 180 may store the future value of acorresponding disk block 178, following, for example, the processing ofa write request in system memory 152 before a successful write operationto the respective disk block 178. The foregoing interim period may occurwhile waiting to write the drive cache 176 or waiting to write therespective disk platter 174. Because the system memory 152 may not storea memory block 180 for every disk block 178, a memory block 180 maycorrespond to multiple disk blocks 178, determined either statically ordynamically. When a memory block 180 corresponds to another disk block178, the respective memory block 180 is said to be in an “Invalid” statewith respect to the relevant disk block 178. In some embodiments, amemory block 180 may also be known as a memory buffer, meaning that thein-memory data structure buffers read/writes to, for example, hard-diskdrives 150.

There may also be blocks of data in persistent memory 156 thatcorrespond to a particular disk block 178. Data blocks written tohard-disk drives 150 may be journaled in the persistent memory 156.These copies of the data are referred to as journal blocks 190.Typically, there is a single memory block 180 and a single cache block176, at any given time, that correspond to a particular disk block 178.There may be multiple journal blocks 190, however, that correspond to aparticular disk block 178. This is because a journal may have a recordof multiple transactions that write to the same disk block 178, or asingle transaction that writes multiple times to the same block.

Thus, in the illustrated embodiment, the location of a particular diskblock 178 may have a corresponding cache block 176, a correspondingmemory block 180, and multiple corresponding journal blocks 190, thoughthe values of those corresponding data blocks may be different dependingon the state of a particular transaction. Moreover, a data valuecorresponding to a particular transaction may be stored in acorresponding memory block 180, journal block 190, cache block 176, anddisk block 178. Typically, a data value corresponding to a particulartransaction is not stored in multiple journal blocks 190, thoughmultiple journal blocks 190 may correspond to the same disk block 178 towhich the respective data value is directed. In general, the term “datablock,” as used herein, refers to a value for a group of data that maybe stored in a disk block 178. For example, “data block” is used todescribe, with reference to FIG. 2A discussed in greater detail below,the respective values of separate portions of a transaction. Moreover,disk blocks 178, cache blocks 176, memory blocks 180, and journal blocks190, as used herein, generally refer to a location in the correspondingstorage device.

The following provides an example disk write transaction. When a node102 receives a disk-write request, a message module may be executed byprocessor 151 to process the disk-write request. The message module mayidentify a data value to write and a disk block 178 to which to writethe data value. A write module may store the identified data value in amemory block 180 and associate the memory block 180 with the identifieddisk block 178. A write module may also journal the data value in ajournal block 190 and associate the journal block with the disk block178. The persistent memory 156 may have previously stored a journalblock 178 associated with the disk block 178. A write module may thenflush the data value from the memory block 180 to the hard-disk drive150, which writes the data value first to an associated cache block 176until the identified disk block 178 may be written with the data value.

In general, the word module, as used herein, refers to logic embodied inhardware or firmware, or to a collection of software instructions,possibly having entry and exit points, written in a programminglanguage, such as, for example, C or C++. A software module may becompiled and linked into an executable program, installed in a dynamiclink library, or may be written in an interpreted programming languagesuch as, for example, BASIC, Perl, or Python. It will be appreciatedthat software modules may be callable from other modules or fromthemselves, and/or may be invoked in response to detected events orinterrupts. Software instructions may be embedded in firmware, such asan EPROM. It will be further appreciated that hardware modules may becomprised of connected logic units, such as gates and flip-flops, and/ormay be comprised of programmable units, such as programmable gate arraysor processors. The modules described herein are preferably implementedas software modules, but may be represented in hardware or firmware.Moreover, although in some embodiments a module may be separatelycompiled, in other embodiments a module may represent a subset ofinstructions of a separately compiled program, and may not have aninterface available to other logical program units.

In one embodiment, the distributed system 100 may comprise a variety ofcomputer systems such as, for example, a computer, a server, a smartstorage unit, and so forth. In one embodiment, the computer may be ageneral purpose computer using one or more microprocessors, such as, forexample, a Pentium processor, a Pentium II processor, a Pentium Proprocessor, a Pentium IV processor, an x86 processor, an 8051 processor,a MIPS processor, a Power PC processor, a SPARC processor, an Alphaprocessor, and so forth. The computer may run a variety of operatingsystems that perform standard operating system functions such asopening, reading, writing, and closing a file. It is recognized thatother operating systems may be used, such as, for example, Microsoft®Windows® 3.X, Microsoft® Windows 98, Microsoft® Windows® 2000,Microsoft® Windows® NT, Microsoft® Windows® CE, Microsoft® Windows® ME,Palm Pilot OS, Apple® MacOS®, Disk Operating System (DOS), UNIX, IRIX,Solaris, SunOS, FreeBSD, Linux®, or IBM® OS/2® operating systems.

III. Exemplary Global Transaction System

FIGS. 2A and 2B illustrate embodiments of a networked cluster of nodesthat coordinates storage transactions. FIG. 2A illustrates a transaction[T₁] 200 that a global transaction module divides into 12 separate datablocks 202. It will be appreciated by one skilled in the art that thereare many different ways a global transaction module could dividetransactions into separate data blocks. Although in the illustratedembodiment the data blocks 202 appear equally sized, in otherembodiments data may be portioned in other suitable ways. In theillustrated embodiment, a global transaction module assigns the datablocks 202 to one of the three participant nodes 102 of distributedsystem 100. Node-3 [P_(S)] 102 is assigned three data blocks [d₁, d₄,d₇] 202; Node-6 [P₁] 102 is assigned four data blocks [d₂, d₅, d₈, d₁₀]202; and Node 1 [P₂] 102 is assigned five data blocks [d₃, d₆, d₉, d₁₁,d₁₂] 202. In the illustrated embodiment, the data blocks 202 are alsocommunicated to their respective participant nodes 102. It will beappreciated that a global transaction module may assign the data blocks202 to participant nodes 102 in a variety of suitable manners. Withreference to the embodiment illustrated in FIG. 1A, Node-3 102 is thecoordinator node [C] and the initiator node [i], in addition to beingthe shared participant node [P_(S)]. In one embodiment, Node-3 102, asthe initiator node [i], may execute the module that divides transaction[T₁] 200 into separate data blocks 202 and then assigns them to therespective nodes 102, including Node-3 102, as the shared participantnode [P_(S)]. Node-3 102, as the coordinator node [C], may thencoordinate the execution of transaction [T₁] 200 among the respectiveparticipant nodes 102. Thus, in the illustrated embodiment, a singlenode 102, such as Node-3 102, may perform the functions of an Initiator,a Coordinator, and a (Shared) Participant by executing, for example, aninitiator module, a coordinator module, and a (shared) participantmodule.

FIG. 2B illustrates one embodiment of a protocol for coordinating globaltransactions in a distributed system, such as distributed system 100.FIG. 2B illustrates an exemplary timing chart according to oneembodiment of a commit protocol 208 for a transaction involving aninitiator 210 (shown as “i”), a first participant 212 (shown as “p₁”), asecond participant 214 (shown as “p₂”), a shared participant 216 (shownas “p_(s)”) and a coordinator 218 (shown as “c”). As discussed above,the initiator 210 and the coordinator 218 are on the same node. In theexample shown in FIG. 2B, the shared participant 216 is also on the samenode as the initiator 210 and the coordinator 218. The first participant212 and the second participant 214 are located on remote nodes.

During the transaction, the initiator 210 adds the first participant212, the second participant 214, and the shared participant 216 to thetransaction. As it does so, the initiator 210 sends start messages 219(three shown) to the first participant 212, the second participant 214,and the shared participant 216. When the initiator 210 is ready to tryto commit the transaction, the initiator sends “prepare” messages 220(four shown) to the coordinator 218, the first participant 212, thesecond participant 214, and the shared participant 216. In oneembodiment, the coordinator 218 is configured to return a response 220 ato the “prepare” message 220. Since the initiator 210 and thecoordinator 218 are on the same node, the coordinator 218 receives the“prepare” message 220 before the remote participants 212, 214.

The first participant 212, the second participant 214, and the sharedparticipant 216 respectively log the “prepare” messages 220 anddetermine whether they are prepared to commit the transaction. If theycan commit the transaction, the first participant 212, the secondparticipant 214, and the shared participant 216 each send a “prepared”message 222 (three shown) to the coordinator 218. If the coordinator 218receives all of the “prepared” messages 222, the coordinator 218 sends“commit” messages 224 (two shown) to the first participant 212 thesecond participant 214. The coordinator 218 does not send a “commit”message 224 to the shared participant 216.

After receiving the “commit” messages 224 from the coordinator 218, thefirst participant 212 and the second participant 214 each log theirrespective “commits” and send “committed” messages 226 to the sharedparticipant 216. Thus, the shared participant 216 learns of thetransaction's outcome from the other participants 212, 214. Aftercommitting to the transaction, the first participant 212, the secondparticipant 214 and the shared participant 218 send “committed” messages228 (three shown) to the initiator 210. For garbage collection purposes,the initiator 210 responds by sending “committed” messages 430 to thefirst participant 212, the second participant 214, and the sharedparticipant 216. After receiving the “committed” message 430 from theinitiator 210, the first participant 212, the second participant 214,and the shared participant 216 clear their respective logs and thecommit protocol 208 ends.

The exemplary timing chart shown in FIG. 2B illustrates the commitprotocol 208 when no failures occur. Since the remote participants 212,214 notify the shared participant 216 of the transaction's outcome, theremote participants 212, 214 can resolve the transaction if they bothbecome disconnected from the coordinator 218.

One example of a global transaction system, in which embodiments ofsystems and methods described herein may be implemented, is described inU.S. patent application Ser. No. 11/449,153 entitled “Non-BlockingCommit Protocol Systems and Methods,” filed Jun. 8, 2006, which is acontinuation of U.S. patent application Ser. No. 11/262,306 entitled“Non-Blocking Commit Protocol Systems and Methods,” filed Oct. 28, 2005,which claims priority to Application No. 60/623,843, filed Oct. 29,2004, all of which are hereby incorporated by reference herein in theirentirety.

IV. Exemplary Data Structures for Nonlinear Journaling

FIGS. 3A, 3B, and 3C illustrate embodiments of data structures forimplementing a nonlinear journal 300. In the illustrated embodiment, thenonlinear journal 300 is a two-dimensional linked list configured tosupport global transactions, atomic additions/deletions to/from thejournal, and shadow buffers. In one dimension, the nonlinear journal 300is a linked list of metadata representing separate transactions, whichis anchored to a journal super-block. The transactions are stored in thenonlinear journal 300 in the order the transactions are prepared. Inother words, the transactions are stored in the nonlinear journal 300when a global transaction module sends a “prepared” message, indicatingthat a transaction has been written to persistent memory 156 and may nowbe stored into the journal with a single atomic write. Thus, aparticular transaction may be stored in persistent memory 156 before itis stored (or linked) to the nonlinear journal 300.

The nonlinear journal also includes a linked list of metadatarepresenting the data blocks that correspond to the same transaction. Inthe illustrated embodiment, the data blocks are stored unordered in thenonlinear journal 300, and these data blocks—or “journal”blocks—correspond to unique disk addresses relative to that transaction.The data blocks are stored to persistent memory 156 as isolated journalblocks 190 before they are linked with the other journal blocks 190corresponding to the same transaction. Metadata structures called“descriptors” organize the journal blocks 190 according to theirrespective transactions. The journal blocks and their associateddescriptors are then linked into the journal with a single atomic write.

FIG. 3A illustrates one embodiment of a nonlinear journal 300 thatincludes a journal descriptor 302, transaction descriptors 304, blockdescriptors 306, and journal blocks 190. The journal descriptor 302 isthe global super-block for the nonlinear journal 300, providing areference point for a system module 157 to locate the nonlinear journal300 in persistent memory 156. In one embodiment, the journal descriptor302 may be written in statically allocated memory at a known location inpersistent memory 156. It will be appreciated that there are many waysto implement a journal super-block. In the illustrated embodiment,journal descriptor 302 includes two data fields: next_seq and txn_list.

In the illustrated embodiment, the nonlinear journal 300 is organizedinto separate transactions, which facilitates, for example, a writemodule to guarantee that either all data blocks in a transaction arewritten to disk or no data blocks in a transaction are written to disk.In the illustrated embodiment, each transaction descriptor 304corresponds to a particular transaction, such as transaction [T₁] 200.The transaction descriptors 304 are organized into a list, thetransaction list 308, which is linked to the journal descriptor 302. Inthe illustrated embodiment, transaction descriptors 304 include metadatafor locating the block descriptors 304 that correspond to the respectivetransaction, for locating the next transaction descriptor 306 in thetransaction list 308, and for managing the respective transaction,including support for global transactions, in the distributed system100. Transaction descriptors 304 include, in the illustrated embodiment,the following data fields: txn_state, desc_list, txn_link,num_participants, and participants. It will be appreciated that thereare many suitable ways to implement transaction descriptors. In otherembodiments, transaction descriptors may include data fields included bytransaction descriptors 304, other data fields, combinations of thesame, and their suitable equivalents. Furthermore, a linked list oftransaction descriptors, such as transaction list 308, may be organizedinto many suitable data structures, including hierarchical datastructures, such as a linked list, a linked list of linked lists, alinked list of multiple dimension linked lists, other data structures,such as a binary tree, a b-tree, and so forth, suitable combinationsand/or equivalents. Moreover, it will be appreciated by one skilled inthe art that transaction descriptors 304 may also be implemented instatically allocated data structures, such as arrays, and suitablecombinations of statically and dynamically allocated data structures.

In the illustrated embodiment, the data blocks 202 of a transaction 200are recorded in respective journal blocks 190 in the nonlinear journal300. The journal blocks 190 are connected indirectly to a transactiondescriptor 304. The respective transaction descriptor 304 corresponds tothe transaction whose data blocks correspond to the respective journalblocks 190. Journal blocks 190 are connected indirectly to respectivetransaction descriptors through block descriptors 306. In theillustrated embodiment, the block descriptors 306 corresponding to aparticular transaction are linked together in a list, the block list310, which is linked to the respective transaction descriptor 304 forthe particular transaction. Block descriptors 306 include metadata forlocating a set of journal blocks 190, for locating the respective diskblocks 178 to which the members of the set of journal blocks 190correspond, and for locating the next block descriptor 306 in the blocklist 310. In the illustrated embodiment, block descriptors 306 includethe following data fields: desc_link, drive, and disk_block. In theillustrated embodiment, block descriptors 306 may include metadata for asingle journal block 190 or multiple journal blocks 190. It will beappreciated by one skilled in the art that other embodiments of blockdescriptors may be implemented in many suitable implementations toinclude metadata for different numbers of journal blocks 190. It will befurther appreciated that other embodiments of block descriptors mayinclude the foregoing data fields, other suitable data fields, suitablecombinations, and suitable equivalents.

Although the illustrated embodiment of nonlinear journal 300 includes ajournal descriptor 302, multiple transaction descriptors 304, andmultiple block descriptors 306, it will be appreciated that there aremany suitable data structures for implementing a nonlinear journal. Forexample, a nonlinear journal may include a statically allocated array oftransaction descriptors that link to linked lists of block descriptors.Additionally and/or alternatively, a particular block descriptor maystore a journal block as a static data field rather than storing a linkto the respective journal block. Additionally and/or alternatively, anonlinear journal may include other descriptors for managing data storedin a nonlinear journal and for performing the desired functions of anonlinear journal. As used herein, a nonlinear journal may describe manysuitable combinations of static and/or dynamic memory allocation,different data structures, and various data fields. In general, the term“nonlinear journal” is used to distinguish ring-based (or linear)journals that overwrite statically allocated memory in a linear fashion.

FIG. 3B illustrates one embodiment of a nonlinear journal that includesa block list 310 corresponding to transaction [T₁] 200, described abovewith reference to FIG. 2A. In the illustrated embodiment, the nonlinearjournal 300 is the journal for Node-3 [P_(S)] 102 in distributed system100. As described above with reference to FIG. 2A, a global transactionmodule assigned three data blocks [d₁, d₄, d₇] 202 to Node-3 [P_(S)]102.

In the illustrated embodiment, journal descriptor 302 includes two datafields: next_seq data field 352 and txn_list data field 354. Thenext_seq data field 352 stores the next sequence number available to anew transaction 200. Transactions 200 may be assigned a unique sequencenumber in order to identify the transactions 200. In general, sequencenumbers remain unique in the face of reboots and power failures, so therecord of the next sequence number is preferably stored in persistentmemory, such as persistent memory 152. In one embodiment, the next_seqdata field 352 holds the next available sequence number. A range ofsequence numbers may be allocated by incrementing the next_seq datafield 352 by the amount of the range. Allocating multiple sequencenumbers at the same time may reduce writes to persistent storage 152. Insome embodiments, the transaction descriptors 304 and their associatedother descriptors and journal blocks 190 may store the unique sequencenumber of the corresponding transaction.

The txn_list data field 354 stores the location in persistent memory 156of the first transaction descriptor 304. In the illustrated embodiment,transaction descriptor [T₁] 304 corresponds to transaction [T₁] 200,described above with reference to FIG. 2A. Although not illustrated,transaction list 308 may include other transaction descriptors 304 withtheir associated block descriptors 306 and journal blocks 190.

In the illustrated embodiment, a transaction descriptor [T₁] 304includes five data fields: txn_state data field 356, desc_list datafield 358, txn_link data field 360, num_participants data field 362, andparticipants data field 364. In the illustrated embodiment, thetxn_state data field 356 stores the value “committed,” indicating thepresent state of the transaction. In one embodiment, the value of thetxn_state data field 356 may correspond to a state defined by a globaltransaction module, such as the global transaction module describedabove with reference to FIG. 2B. Additionally and/or alternatively, thevalue may correspond to a state defined by a local transaction module.Transaction states for both global and local transactions are describedin more detail below with respect to FIGS. 4A and 4B, respectively. Inthe illustrated embodiment, the desc_list data field 358 stores thelocation in persistent memory 156 for the head of a linked list of blockdescriptors [B₂, B₁] 306, which link to the journal blocks 190 thatcorrespond to transaction [T₁] 200. In the illustrated embodiment, thetxn_link data field 360 stores a location in persistent memory 156corresponding to the next transaction descriptor 304 in the transactionlist 308. If transaction descriptor [T₁] 304 is the last transactiondescriptor 304 in the transaction list 308, then the respective txn_linkdata field 360 stores a null pointer—that is a value that indicates noaccessible memory address location is being stored. In the illustratedembodiment, the num_participants data field 362 stores the number “3,”indicating that there are three participants for transaction [T₁] 200.In the illustrated embodiment, the participants data field 364 stores alist of identifiers “1, 3, 6,” which identify the participant nodes(Node-1, Node-3, and Node-6) 102 for transaction [T₁] 200.

In the illustrated embodiment, block descriptors [B₂, B₁] 306 include aconstant data field, desc_link data field 366, and multiple groups ofthree data fields: drive data field 368, disk_block data field 370, anddata_link data field 372. In the illustrated embodiment, the desc_linkdata field 366 of block descriptor [B₂] stores the location inpersistent memory 156 for block descriptor [B₁] 306, which is the nextblock descriptor 306 in the block list 310. Because block descriptor[B₁] 306 is the last block descriptor 306 in the block list 310, itsdesc_link data field 366 stores a null pointer, indicating that thereare no remaining block descriptors 306 in the block list 310. In theillustrated embodiment, data_link data field 372 of block descriptor[B₂] 306 stores a location for the journal block 190 that stores thevalue of data block [d₇] 202 of transaction [T₁] 200. In the illustratedembodiment, the drive data field 368 and the disk_block data field 370,of block descriptor [B₂] 306, store the values “7” and “18,”respectively, indicating that the value of data block [d₇] 202 oftransaction [T₁] 200 is assigned for storage to disk block [18] 178 ofhard-disk drive [7] 150 on Node-3 102 of distributed system 100.Respective data fields for block descriptor [B₁] 306 indicate that thevalue of data block [d₄] 202 of transaction [T₁] 200 is assigned forstorage to disk block [43] 178 of hard-disk drive [4] 150 on Node-3 102of distributed system 100, and that the value of data block [d₁] 202 oftransaction [T₁] 200 is assigned for storage to disk block [21] 178 ofhard-disk drive [5] 150 on Node-3 102 of distributed system 100. Inother embodiments, the value of the respective data block 202 may bestored in place of the data_link data field 372.

FIG. 3C illustrates one embodiment of defragmenting a nonlinear journal300 by deleting a block descriptor 306 and linking its journal blocks190 to another block descriptor 306. In some embodiments, there may be aneed to implement defragmentation because there may be heterogeneousblock sizes allocated in the nonlinear journal 300 and because there maybe different lifetimes of data. In one embodiment, the basic unit ofallocation may be made to be 512 bytes, and larger blocks may beconstructed with scatter/gather lists. In another embodiment, thenonlinear journal 300 may be defragmented on demand. In the illustratedembodiment, to move the data of any particular journal block 190 and/orits associated descriptors in the nonlinear journal 300, the data ordescriptor may be copied to a new location and a single atomic write maychange the pointer in the nonlinear journal 300 that points to therespective data or descriptor. Additionally and/or alternatively, blockdescriptors 306 may be coalesced in the case that they are adjacent inthe block list 310. A single atomic write may be used to change anylinkage. For example, a single block descriptor 306 may be written withlinks to the combined set of journal blocks 190 corresponding to twodifferent block descriptors 306. In the illustrated embodiment, a newblock descriptor 306 is first created with its respective data_link datafields 372 linked to the journal blocks 190 of the two old blockdescriptors 306. The new block descriptor 306 is then linked into therespective block list 310 in place of the old block descriptors 306. Thetwo old block descriptors 306 may now be freed.

In other embodiments, there may be transaction descriptors, whichdirectly reference the journal blocks without block descriptors. In oneembodiment, the transaction descriptors could be variable-sized inpersistent memory. In another embodiment, there may be blockdescriptors, which duplicate transaction metadata in the blockdescriptors themselves without implementing transaction descriptors. Inone embodiment, a scheme may be used to determine which blockdescriptor(s) has the correct transaction state. In another embodiment,transaction descriptors are pre-allocated in a contiguous array. In oneembodiment, block descriptors are also pre-allocated in a contiguousarray. In another embodiment, transaction descriptors share blockdescriptors. In one embodiment, the transaction descriptors usereference counts. In another embodiment, block descriptors include datainline. In another embodiment, partial block descriptors may be placedinside transaction descriptors, for example, to save space. Theembodiments enumerated above are by way of example to illustrate themany suitable organizations and configurations comprised in embodimentsof the invention. Other embodiments may specify other descriptors, datafields, and structures. Still other embodiments may combine theabove-enumerated embodiments and their suitable equivalents.

V. Exemplary Global Transactions

FIGS. 4A and 4B illustrate state diagrams of embodiments of a nonlinearjournal system for global and local transactions, respectively. FIG. 4Ais a state diagram showing the on-journal states for one embodiment of aglobal transaction module using one embodiment of a nonlinear journal300. State U is the unknown state. In state U, there is no knowledge ofthe transaction in the nonlinear journal 300. State W is the writingstate. In the writing state, data blocks 202 are written to persistentmemory 156 as journal blocks 190, and the linked list of blockdescriptors 306 is built up. In one embodiment, transactions in state Uand W will not be reconstructed at replay time because they are not yetlinked into the transaction list 308. In other words, if the nonlinearjournal 300 is replayed after a failure, the data written to persistentmemory 156 for all transactions in state U will be ignored because thereplay module will only restore the data linked into thetransaction-list 308. State P is the prepared state. A preparedtransaction has a transaction descriptor 304 that is linked into thetransaction list 308, and its descriptors' txn_state data field 356 isset to “prepared.” State C is the committed state. A committedtransaction is linked into the transaction list 308, and itsdescriptor's txn_state data field 356 is set to “committed.” A committedtransaction is waiting for the global transaction module to confirm thatthe participants have committed before transitioning to done. The finalstate D is the done state. A transaction in the done state holds thejournal blocks 190 that still need to be written to disk.

The following are exemplary entry points into one embodiment of ajournal subsystem for a participant in a distributed system, such as aparticipant node 102 of distributed system 100:

-   -   start( ): Called to start a transaction.    -   write(b): Called to write a block on a transaction.    -   delta(b, o, d): Called to write a delta to a block on a        transaction.    -   prepare(P, S): Called at prepare time. It is passed the set of        participants P and the shared participant S which gets put in        the txn descriptor.    -   commit( ): Called to commit the transaction.    -   abort( ): Called to abort the transaction.    -   done( ): Called to indicate it is no longer necessary for the        journal to keep a record of the transaction's outcome.    -   block_synced(b): Called when the block is known to be on the        disk platters.

The following is exemplary pseudocode of one embodiment of a journalsubsystem in a participant node 102 of distributed system 100: (In theillustrated embodiment, the pseudocode is common to both global andlocal transactions of the journal subsystem.)

function abort_blocks(B): for all b in B:  block_abort(b) functionflush_blocks(B): for all b in B:  block_flush (b) function link(P, S,s): write out txn descriptor with participants P, shared participant Sand state s do atomic write to link descriptor into the global txn listfunction unlink( ): do atomic write to unlink descriptor from the globaltxn list set state to U function link_data(b): do atomic write to linkdata block into txn block list function unlink_data(b): do atomic writeto unlink data block from txn block list

The following exemplary pseudocode further describes one embodiment of ajournal subsystem in a participant node 102 of distributed system 100.In the illustrated embodiment, the pseudocode is with respect to globaltransactions:

Function mark_state(s): do atomic write to change state to s in state U:in txn descriptor  on start( ): set state to (W, Ø) in state (W, B):  onwrite(b): block_alloc( ) link_data(b) set state to (W, B U {b})  ondelta(b, o, d): If (predecessor_block(b))  set state to (W, D U {b, o,d}) else  read(b)  Apply_delta(b, o, d)  write(b)  on prepare (P, S):Link(P, S, PREPARED) set state to (P, P, S, B) on abort( ):abort_blocks(B) set state to U in state (P, P, S, B):  on commit ( ):mark_state(COMMITTED) flush_blocks(B) set state to (C, P, B)  on abort(): abort_blocks(B) unlink( ) In state (C, P, B):  on block_synced(b):unlink_data(b) set state to (C, P, B \ {b})  on done( ) and B = Ø:unlink( )  on done( ): mark_state(DONE) set state to (D, B) in state (D,B):  on block_synced(b): unlink_data(b) set state to (D, B \ {b})  on B= Ø: unlink( )

FIG. 4B is the state diagram showing the on-journal states for oneembodiment of a local transaction module using one embodiment of anonlinear journal 300. Local transactions are those transactions thataffect relevant node 102 executing the transaction, but do not affectother nodes 102 in the distributed system. Because there is no need tohold onto the outcome of the transaction for other participants toquery, a local transaction goes directly from W to D.

The following exemplary pseudocode further describes nonlinearjournaling with respect to local transactions:

In state U:  on start( ): set state to (W,Ø) in state (W, B):  onwrite(b): block_alloc( ) link_data(b) set state to(W, B U {b})  ondelta(b, o, d): if (predecessor_block(b))  set state to (W, D U {b, o,d}) else  read(b)  apply_delta(b,o,d)  write(b)  on commit( ): link(Ø,Ø,DONE) flush_blocks(B) set state to (D, B)  on abort( ): abort_blocks(B)set state to U in state (D, B):  on block_synced(b): unlink_data(b) setstate to (V, B \ {b})  on B = Ø: unlink( )

FIGS. 4C and 4D illustrate state diagrams describing the life cycle ofjournal blocks 190 in nonlinear journal 300. FIG. 4E illustrates a statediagram describing one embodiment of providing support for down drivesin nonlinear journal 300. These state diagrams are described in greaterdetail below in Sections VI and VIII, respectively.

A. Example Transactions

FIGS. 5A, 5B, and 5C illustrate embodiments of building, modifying, andremoving nonlinear journal data structures. Up until the point atransaction is prepared (or done for local transactions), in theillustrated embodiment, the transaction is guaranteed to be ignored (toabort) at replay time because a transaction is not linked into thetransaction list 308 until the transaction is prepared. Accordingly, inthe illustrated embodiment, a list of journal blocks 190—or, morespecifically, a list of block descriptors 306 with their associatedjournal blocks 306—may be constructed in any free space in persistentmemory 156, unattached to any other structures.

In the illustrated embodiment, to construct the block list 310, thefirst block descriptor 306 is written with a null desc_link data field366. Subsequent block descriptors 306 are written with a desc_link datafield 366 pointing to the previous block descriptor 306. With referenceto FIGS. 3A and 3B, in one embodiment, the block descriptors 306 arewritten from right to left. In the illustrated embodiment, to replayproperly, either the journal blocks 190 need to be replayed in the sameorder as they were written, or there needs to be no duplicate blockwrites to the same disk location in a single transaction. In theillustrated embodiment, there are no journal blocks 190 in the sameblock list 310 that correspond to the same disk block 178 on the samehard-disk drive 150. In other embodiments, the journal blocks may bereplayed in order, and journal blocks 190 may correspond to the samedisk location within the same transaction. In one embodiment, when adisk block 178 is written by multiple data blocks 202 within the sametransaction, a single corresponding journal block 190 may be overwrittenwith the new data.

Journal-block writes to persistent memory 156 may be asynchronous. Toprepare a transaction, the journal-block writes are completed beforelinking the corresponding transaction descriptor 304 for the respectivejournal blocks 190 into the transaction list 308. Because a transactionis not prepared until after it is linked, the transaction-descriptorwrite may be asynchronous as well. After the block-descriptor writes arestarted, the respective transaction descriptor 304 is written with itsdesc_list data field 358 set to the location of a block descriptor 306,its txn_link data field 360 set as a null pointer, and its txn_statedata field 356 set to “prepared.” When the descriptor writes andjournal-block writes finish, a single atomic write to the txn_link datafield 360 of the previous descriptor in the transaction list 308 linksthe new transaction descriptor 304 and its associated block descriptors306 and respective journal blocks 190.

To commit a global transaction, a single atomic write is made to thetxn_state data field 356 of the respective transaction descriptor 304.That field is set to “committed.” For a local transaction, the txn_statedata field 356 of the respective transaction descriptor 304 is writtenas “done” when the transaction descriptor 304 is written. Thetransaction descriptor 304 is then atomically linked into thetransaction list 308.

In one embodiment, before removing a record of a global transaction fromthe nonlinear journal 300, the contents of respective journal blocks 190associated with the transaction are written to corresponding disk blocks178, and a global transaction module, such as the coordinator moduledescribed above with reference to FIG. 2B, informs the respectiveparticipant node 102 that the remaining participant nodes 102 arecommitted. Local transactions, however, are done as soon as they commit.Local transactions are described in greater detail above with referenceto FIG. 4B. Once journal blocks 190 are stored to corresponding diskblocks 178, journal blocks 190 associated with a particular transactioncan be freed while that transaction is still active. This may be done byzeroing or nulling the data_link data field 372 of the associated blockdescriptor 306 with an atomic write. If the associated block descriptor306 no longer links to any journal blocks 190, it may be removed byunlinking it from the desc_link data field 366 of the preceding blockdescriptor 306, or the desc_list data field 358 of the associatedtransaction descriptor 304 (if the block descriptor 360 was the firstblock descriptor 306 in block list 310). Any remaining block descriptors306 in the block list 310 may then be relinked to the respectivedesc_link data field 366 or desc_list data field 358, which was just setto null to unlink it from the removed block descriptor 306. Once journalblocks 190 associated with a transaction are stored to correspondingdisk blocks 178 and freed from the nonlinear journal 300, thecorresponding transaction descriptor 304 is a candidate for being freedas well. It may be retained, however, until the global transactionmodule releases it, for example, by setting the state of the transactionto done.

FIG. 5A illustrates one embodiment of building transaction [T₁] 200 inthe nonlinear journal 300. In state 500, Node-3 [P_(S)] 102 ofdistributed system 100 receives a “start T₁” message from an initiatornode, such as the initiator node [i] 102 described above with referenceto FIG. 2B. In the illustrated embodiment, prior to receiving the “startT₁” message, journal descriptor 302 was written to persistent memory 156in a known location. In the illustrated embodiment, there are notransaction descriptors 304 linked in the transaction list 308,indicating that no transaction is currently journaled in the nonlinearjournal 300. Upon receiving the “start” message, the participant modulereserves space for transaction descriptor [T₁] 304 in space 504 inpersistent storage 156.

Throughout the description of the drawings reference is made to a“participant module.” The participant module is meant to refer tosuitable programs implemented in hardware or software that execute thedescribed functions and operations. In some embodiments, the participantmodule may be a journal subsystem that is responsible for managing theoperations associated with a journal, such as nonlinear journal 300. Inother embodiments, the participant module may comprise a journalsubsystem along with other subsystems, such as a subsystem, module,procedure, or process for global transactions. In still otherembodiments, the participant module may be a component of a journalsubsystem. The participant module need not be separately compilable,though it may be. In one embodiment the participant module may executeon a system comprising a single node, such as node 102, executing ajournal subsystem in isolation from other entities. In otherembodiments, the participant module may execute on a distributed system,such as distributed system 100, executing in separate instances onindividual nodes, or possibly executing in part on multiple nodes withinthe system.

In 506, the participant module receives a “write [d₁]” message from aninitiator node, such as the initiator node [i] 102 described above withreference to FIG. 2B. The participant module reserves space for blockdescriptor [Be] 306 in space 506 of persistent memory 156. Theparticipant module then writes a journal block 190 with the value ofdata block [d₁] 202. Because the block descriptors 306 in theillustrated embodiment may link to two journal blocks 190, blockdescriptor [Be] 306 is not yet written to persistent storage 156.

In 512, the participant module receives a “write [d₄]” message from aninitiator node, such as the initiator node [i] 102 described above withreference to FIG. 2B. The participant module writes a journal block 190with the value of data block [d₄] 202.

In 518, the participant module receives a “write [d₇]” message from aninitiator node, such as the initiator node [i] 102 described above withreference to FIG. 2B. The participant module writes block descriptor[B₁] 306, links to it journal block [d₁] 190 and journal block [d₄] 190,and writes their respective drive data fields 368 and disk_block datafields 370. The participant module reserves space for block descriptor[B₂] 306 in space 520 of persistent memory 156, and writes journal block[d₇] 190 in persistent memory 156 with the value of data block [d₇] 202.Although in the illustrated embodiment the respective block descriptor306 is not written until a second journal block 190 is written, in otherembodiments a block descriptor 306 may be written after (or even before)a first journal block 190 is written. It will be appreciated by oneskilled in the art that there are many suitable ways to order thewriting of journal blocks 190 and their associated descriptors.Moreover, although in the illustrated embodiment the block descriptors306 comprise metadata for two journal blocks 190, block descriptors maybe designed to comprise many suitable numbers of journal blocks 190.

In 524, the participant module receives a “prepare (T₁)” message from aninitiator node, such as the initiator node [i] 102 described above withreference to FIG. 2B. The participant module writes block descriptor[B₂] to persistent storage 156, links to it block descriptor [B₁] andjournal block [d₇] 190, and stores values for drive data field 368 anddisk_block data field 370 corresponding to data block [d₇] 202. Theparticipant module then writes transaction descriptor [T₁] 304, whichlinks to it block descriptor [B₂] 306 and keeps respective values forits txn_state data field 356 (set to “prepared”), txn_link data field360, num_participants data field 362, and participants data field 364.The participant module then links transaction descriptor [T₁] 304 intothe transaction list 308. Because transaction descriptor [T₁] 304 is thefirst transaction descriptor to be linked into the transaction list 308,the participant module writes to the txn_list field 354 of journaldescriptor 302 the location of transaction descriptor [T₁] 304 inpersistent memory 156.

In 530, the participant module receives a “commit (T₁)” message from acoordinator node, such as the coordinator node [C] 102 described abovewith reference to FIG. 2B. As described above with reference to FIG. 2B,in some embodiments of a global transaction module the coordinator nodesdoes not send a “commit” message, waiting instead for the otherparticipant nodes to inform the shared participant that the transactionis committed. In the illustrated embodiment, however, the coordinatornode [C] 102 sends a “commit (T₁)” message. To commit the transaction,the participant module atomically writes to the txn_state field 356 oftransaction descriptor [T₁] 304 the value “committed.”

FIG. 5B illustrates the effect of an “abort (T₁)” message received after518. In 536, the participant module receives the message following state518. The participant module frees the space allocated to journal blocks[d₁, d₄, d₇] 190, block descriptor [B₁] 306, and transaction descriptor[T₁] 304. There is nothing to unlink from nonlinear journal 300 becausetransaction [T₁] 200 had not yet been prepared. If the nonlinear journal300 had been replayed prior to the “abort T₁,” the journal blocks [d₁,d₄, d₇] 190 would not have been reconstructed because transaction [T₁]200 had never been linked into nonlinear journal 300.

FIG. 5C illustrates the effect of an “abort (T₁)” message received after524. In 542, the participant module receives the message following 524.The participant module unlinks transaction descriptor [T₁] 304 from thetransaction list 308 by atomically writing the txn_list data field 354of journal descriptor 302 and setting it to null. The participant modulethen frees the space allocated to journal blocks [d₁, d₄, d₇] 190, blockdescriptors [B₁, B₂] 306, and transaction descriptor [T₁] 304. In theillustrated embodiment, after unlinking the transaction and freeing thespace allocated with its descriptors, the nonlinear journal 300 stillincludes the journal descriptor 302. Although in the illustratedembodiment the participant module immediately freed the space allocatedto the journal blocks 190 and their associated descriptors, in otherembodiments the space may not be freed until a garbage collector isexecuted. In still other embodiments, the space may not be freed untilthe space is needed for another journal block 190 or descriptor. Forexample, a space allocator may determine whether allocated journal spaceis active by searching through in-memory data structures for a referenceto the space and reclaim any unreferenced space. Even a journal block190 that has not been linked to by a descriptor could still bereferenced in memory, allowing the space allocator to distinguishbetween unlinked and leftover journal blocks 190.

FIG. 5D, discussed in greater detail below in Section VI(C), illustratesthe freeing of space allocated for a particular transaction 200 asrespective journal blocks 190 are synced with their corresponding diskblocks 178 following state 530.

B. Example Procedures

FIGS. 6A and 6B illustrate flowcharts of embodiments of building,modifying, and removing nonlinear journal data structures. FIG. 6Aillustrates one embodiment of building a block list 310 corresponding toa particular transaction 200 following a “write” message. In state 602,upon receiving a “write” message, the participant module determineswhether there is a block descriptor 306 with space available for anadditional journal block 190. If there is a block descriptor withavailable space, the participant module proceeds to state 610. If thereis no block descriptor 306 with available space, the participant moduleproceeds to state 604. In state 604, the participant module determineswhether there is a block descriptor 306 that has not been written topersistent storage 156. If there is a block descriptor 306 that has notbeen written, then the participant module proceeds to state 606. If itis not the case that there is a block descriptor 306 that has not beenwritten, then the participant module proceeds to state 608. In state606, the participant module writes the unwritten block descriptor 306,and then proceeds to state 608. Writing a block descriptor 306 isdescribed above in further detail with respect to FIG. 3B. In state 608,the participant module reserves space for a new block descriptor 306. Instate 610, the participant module writes the relevant journal block 190.Although in the illustrated embodiment an allocated but unwritten blockdescriptor 306 is written before a new block descriptor 306 is written,in other embodiments block descriptors 306 for a particular transactionmay wait to be written until after the journal blocks 190 for thetransaction have been written. It will be appreciated by one skilled inthe art that there are many suitable ways to order the writing ofjournal blocks 190 and their associated descriptors. For example, in oneembodiment block descriptors 306 may include metadata for a singlejournal block 190 and be written before or after the respective journalblock 190 is written.

FIG. 6B illustrates a flow chart of one embodiment of building atransaction list 308 for a nonlinear journal 300 following a “prepare”message. Because the most recently written block descriptor 306 is notwritten, in the illustrated embodiment, following the most recent“write” message, there is an unwritten block descriptor 306 inpersistent memory 156 when the “prepare” message is received. In state642, the participant module writes the unwritten block descriptor 306 topersistent memory 156. As described above with reference to FIG. 6A, insome embodiments block descriptors 306 may be written at varioussuitable moments; hence, there may not be an unwritten block descriptor306. In some embodiments, the participant module may determine whetherthere is an unwritten block descriptor 306. In other embodiments, theparticipant module may know whether there is an unwritten blockdescriptor 306 based on the specific implementation. In state 644, theparticipant module writes the unwritten transaction descriptor 304 topersistent memory 156 with its txn_state data field 356 set to“prepared.” Writing a transaction descriptor 304 is discussed above infurther detail with respect to FIG. 3B. In state 648, the participantmodule links the relevant transaction descriptor 304 to the transactionlist 308.

VI. Exemplary Journal Space Management

FIGS. 4C and 4D illustrate state diagrams describing the life cycle ofjournal blocks 190 in nonlinear journal 300. In the illustratedembodiment, journal blocks 190 and their associated descriptors are heldin nonlinear journal 300 until the corresponding disk blocks 178 havebeen written. In one embodiment, the participant module instructs therespective hard-disk drive 150, or possibly a subset of hard-disk drives150, to flush the contents of its drive cache 172, causing the contentsof the cache blocks 176 to be written to the respective disk blocks 178.The term subset, as used herein, is given its ordinary mathematicalmeaning, which includes the possibility that the entire set is a subsetof a set. Additionally and/or alternatively, the hard-disk drives 150may be self-flushing, regularly flushing the contents of theirrespective drive cache 172. In some embodiments, the hard-disk drivesmay communicate to a participant module when their respective drivecache 172 has been flushed. Additionally and/or alternatively, theparticipant module may be programmed to expect such regular flushes.

In still other embodiments, journal blocks 190 are released on a leastrecently used (LRU) basis. Some of these alternative embodiments arediscussed in greater detail below following the more detailed discussionof the cache-flush (or synchronization) embodiments, which followsimmediately.

A. Sync Expiration

In the illustrated embodiment, journal blocks 190 start in theunallocated state. U. They enter the pinned state P when they areallocated. Once they enter state P, they are held in the nonlinearjournal 300 and are only released when they are flushed or aborted.Journal blocks 190 in the flushed state F have been written to theircorresponding hard-disk drives 150 and are now sitting in anindeterminate state, possibly in the drive cache 172 of the respectivehard-disk drive 150, waiting to be written to the respective diskplatter 174. In the illustrated embodiment, a list of such blocks D_(F)is retained until the next drive sync completes, at which time the listis detached and cleared, and all such blocks are sent back to theunallocated state U. In one embodiment, drive sync is implemented as abackground process (or thread) that may periodically issue the syncs,and may be woken up if space is needed.

In one embodiment, journal space is managed in system memory, such assystem memory 152, with allocate and free logic. Journal space may bemanaged in system memory such that free blocks may be found as necessaryand marked as used, and used blocks may be marked as free when no longerreferenced. In other embodiments, journal space may be managedexternally in persistent memory, such as persistent memory 156, throughuse of a table or a linked list of free blocks.

The following are exemplary entry points into one embodiment of a blockhandling subsystem for a nonlinear journal, such as nonlinear journal300:

-   -   block_alloc( ): Allocate a block in the journal    -   block_abort( ): Free a block immediately    -   block_flush( ):Flush contents of journal block to disk and free        it from the journal when it is safely on the disk

The following exemplary pseudocode further describes space allocationand journal block lifecycle in one embodiment of a block handlingsubsystem for a nonlinear journal, such as nonlinear journal 300:

In state U:  on block_alloc( ): Dp = Dp U {b} set state to P in state P: on block_abort( ): Dp = Dp \ {b} set state to U  on block_flush( ):write block to disk Dp=Dp\{b} D_(P) = D_(P) U {b} set state to F instate F:  on block_synced( ): set state to U  on block_lru( ): set stateto U

In one embodiment, the values of data blocks 202 are flushed to theirrespective hard-disk drives 150 indirectly from the journal blocks 190of nonlinear journal 300 in persistent storage 156. In anotherembodiment, the values of data blocks 202 are flushed to theirrespective hard-disk drives 150 directly from the memory blocks 180 insystem memory 152. Indirectly flushing from the journal blocks 190 maybe advantageous because it frees system memory 152 for reuse. Directlyflushing from the memory blocks 180, however, may be faster.

In one embodiment, the block abort function in the pseudocode aboverefers to marking a block as being synced (and, thus, available forexpiration from a journal), and then freeing it immediately, rather thanwaiting for garbage collection. It will be appreciated by one skilled inthe art that many suitable functions may supply the functionalitydescribed in the pseudocode above.

The following exemplary pseudocode further describes one embodiment ofdisk synchronization for a nonlinear journal, such as nonlinear journal300:

function drive_synced(B):  for all b in B:   Block_synced(b) functionstart_syncer(B):  send sync command to drives  set state to (S, B,) instate (W, D_(F)):  on Need Space: Start_syncer(D_(F))  on D_(F) >D_(Fmax): Start_syncer(D_(F)) in state (S, B, D_(F)):  on sync return:drive_synced(B) Set state to (W, DF)

In one embodiment, when allocating data blocks 190, if space is notavailable, the nonlinear journal 300 may block waiting for journalblocks 190 to be flushed or aborted. If the journal blocks 190 areassociated with transactions in the writing state or associated withdown drives, in one embodiment, this may lead to deadlock and may bedealt with, for example, by failing transactions. Down drives are drivesthat are not currently available, but that may reenter the system; theyare discussed in greater detail below.

B. Alternative LRU Expiration

In one embodiment, a least recently used (LRU) module for journal-blockexpiration may be used. This model attempts to ensure that data iswritten through the drive caches 172 to the disk platters 174 by storingas much data as possible in persistent storage 156 and rewriting all ofthe data on replay. Replaying data is discussed in further detail belowwith reference to FIG. 7. Data associated with flushed transactions arefreed from the nonlinear journal 300 in this model in LRU order, as newtransactions allocate data. In one embodiment, the nonlinear journal 300can operate a LRU model in a per-drive fashion. The nonlinear journal300 may reserve a certain minimum number of journal blocks 190 that havebeen written to a respective hard-disk drive 150. A particular minimumreserve may be defined as sufficient to replay over the drive cache 172of a respective hard-disk drive 150, and the LRU module may not expirejournal blocks 190 if it would reduce the number of flushed blocks for arespective hard-disk drive 150 below this limit.

In one embodiment of an LRU implementation, FIGS. 4C and 4D describe theappropriate state diagrams, except that the mechanism to move journalblocks 190 from the F state to the U state is different. In theillustrated embodiment, Journal blocks 190 in the F state are retainedin the nonlinear journal 300 until the number of blocks associated witha respective hard-disk drive 150 exceeds the disk replay minimum,D_(Fmin), or possibly a timeout expires that represents the longest timea cache block 176 can exist in the drive cache 172. In the illustratedembodiment, once a journal block 190 is eligible for LRU, it may beretained until demand causes an LRU flush command, or when theassociated hard-disk drive 150 is kicked out. Journal blocks 190corresponding to hard-disk drives 150 that are down are typically notfreed, but may be LRU'd until the D_(Fmin) limit is reached. Down drivesare discussed in more detail below.

In one embodiment, a nonlinear journal, such as nonlinear journal 300,may support both LRU and drive sync modules simultaneously. In, thishybrid module, once a journal block 190 has been flushed to a respectivehard-disk drive 150, it cannot be released for LRU until either a drivesync or another condition is satisfied. The exit condition may beconfigurable and may include both requirements at once.

C. Descriptor Expiration

Although journal blocks 190 may be expired from the nonlinear journal300 in an LRU or sync fashion, the associated block descriptors 306 andtransaction descriptors 304 may be expired from the nonlinear journal300 as a side effect of their associated journal blocks 190 expiring.When journal blocks 190 are expired from the journal, either throughdrive sync or LRU, their space is released. To prevent a journal block190 from being replayed, the respective data_link data field 372 of thecorresponding block descriptor 306 may be cleared or nulled. This may bedone by atomically rewriting the data_link data field 372 with a nullvalue such as 0. Because block descriptors 306 reference multiplejournal blocks 190, the respective block descriptor 306 may not be freeduntil all journal blocks 190 are unlinked. It will be appreciated thatthere are many ways to free space from the nonlinear journal. To unlinka block descriptor 306, the desc_link field 366 of the previous blockdescriptor 306 may be overwritten with the address for the succeedingblock descriptor 306. Additionally and/or alternatively, the desc_listfield 358 of the relevant transaction descriptor 304 may be overwrittenwith the succeeding address of the succeeding block descriptor 306. Inanother embodiment, block descriptors 304 may remain allocated untiltheir associated transaction descriptors 304 are unlinked from thetransaction list 308. In one embodiment, a transaction descriptor 304may be released when all journal blocks 190 in the block list 310 arereleased and the transaction is done. This is accomplished by removingthe transaction descriptor 304 from the transaction list 308. Because apath to a block descriptor 306 may be through the transaction descriptor304, any allocated block descriptors 306 may be freed directly at thistime as well.

D. Example Sync Expiration

FIG. 5D illustrates the freeing of space allocated for a particulartransaction 200 as respective journal blocks 190 are synced with theircorresponding disk blocks 178 following state 530. In 548, theparticipant module receives a “block_synced (d₄)” message, indicatingthat respective disk block 178 has been synced with journal block [d₄]190. The participant module unlinks journal block [d₄] 190 from blockdescriptor [B₁] 306. The participant module then frees the spaceallocated to journal block [d₄] 190.

In 550, the participant module receives a “done (T₁)” message, whichindicates that all of the participants of transaction [T₁] 200 havecommitted, and the corresponding transaction descriptor [T₁] 304 is nolonger needed in order to rebuild the transaction when necessary. Thetxn_state data field 356 is atomically set to “done.” The transactiondescriptor [T₁] 304, however, is not removed until receiving“block_synced” messages for the journal blocks corresponding to thetransaction descriptor [T₁] 304. Thus, in the illustrated embodiment, atransaction descriptor 304 is removed from the transaction list 308 whenits txn_state data field 356 is set to “done” and when the contents ofits corresponding journal blocks 190 are synchronized with theirrespective disk blocks 178 following the appropriate cache flushes.

In state 554, the participant module receives a “block_synced (d₇)”message. Because journal block [d₇] 190 is the last journal block 190linked to block descriptor (B₂) 306, transaction descriptor (T₁) 304 isrelinked atomically to block descriptor [B₁] 306. The participant modulethen frees the space allocated to journal block [d₇] 190 and blockdescriptor [B₂] 306.

In state 560, the participant module receives a “block_synced (d₁)”message. Because journal block [d₁] 190 is the last remaining journalblock 190 in the block list 310 for transaction descriptor [T₁] 304, andbecause its txn_state data field 356 is set to “done,” transactiondescriptor [T₁] 304 is unlinked from the transaction list 308. Theparticipant module then frees the space allocated to journal block [d₁]190, block descriptor [B₁] 306, and transaction descriptor [T₁] 304,leaving only journal descriptor 302 in the nonlinear journal 300.

Although in the illustrated embodiment the “block_synced” messagessuggest per-block synchronization, these “block_synced” messages may bepart of a synchronization procedure that synchronizes an entirehard-disk drive 150 with an explicit cache flush of the respective drivecache 172. Corresponding “block_synced” messages may then be issued forthe synchronized disk blocks 178 with corresponding journal blocks 190still written to the nonlinear journal 300. It will be appreciated byone skilled in the art that there are many suitable ways to implement asynchronization procedure in accordance with the embodiments describedherein.

E. Example Sync Expiration Procedure

FIG. 6C illustrates a flow chart of one embodiment of removing datablocks and their related descriptors from nonlinear journal 300 and tofree their associated space in persistent memory 156. In state 672, theparticipant module determines whether there is, in addition to therelevant journal block 190, another journal block 190 that is linked tothe relevant block descriptor 306. In the illustrated embodiment, therelevant journal block 190 is the journal block 190 being removedbecause, for example, the corresponding disk block 178 has synced. Inthe illustrated embodiment, the relevant block descriptor 306 is theblock descriptor from which the relevant journal block 190 is beingremoved. If there is another journal block 190, then the participantmodule proceeds to state 686. If there is not another journal block 190,then the participant module proceeds to state 674.

In state 674, the participant module determines whether there is anotherblock descriptor 306 corresponding to the relevant transactiondescriptor 304. The relevant transaction descriptor 304 is thetransaction descriptor 304 from which the relevant journal block 190 isbeing removed. If there is another block descriptor 306 corresponding tothe relevant transaction descriptor 304, then the participant moduleproceeds to state 676. If there is not another block descriptor 306corresponding to the relevant transaction descriptor 304, then theparticipant module proceeds to state 677. In state 677, the participantmodule determines whether the txn_state data field 356 of the relevanttransaction descriptor 304 is set to “done.” If the transaction is done,the participant module proceeds to state 678. If the transaction is notdone, the participant module proceeds to state 676. In state 676, theparticipant module relinks the remaining block descriptors 306 in therelevant block list 310, which unlinks the relevant block descriptor 306from the nonlinear journal 300. In the illustrated embodiment, thisrelinking and unlinking may be performed with a single atomic writeeither to the desc_list data field 358 of the relevant transactiondescriptor 304 (if the removed block descriptor 306 had been linked tothe relevant transaction descriptor 304) or to the desc_link 306 of theblock descriptor 306 that previously linked to the removed blockdescriptor 306. By overwriting either data field, the nonlinear journal300 may simultaneously link and unlink, providing for a consistentjournal state.

In state 678, the participant module determines whether there is anothertransaction descriptor 304 in the transaction list 308. If there isanother transaction descriptor 304 in the transaction list 308, then theparticipant module proceeds to state 680. If there is not anothertransaction descriptor 304 in the transaction list 308, then theparticipant module proceeds to state 682. In state 680, the participantmodule unlinks the relevant transaction descriptor 304 from thetransaction list 308 and relinks the remaining transaction descriptor(s)in the transaction list 308. In the illustrated embodiment, relinkingthe remaining transaction descriptors(s) includes either linking thetxn_list data field 354 of journal descriptor 302 or the txn_link datafield 360 of the preceding transaction descriptor 304 to the transactiondescriptor(s) 304 in transaction list 308 that the relevant transactiondescriptor 304 linked to.

In state 682, the participant module unlinks the relevant transactiondescriptor 304 from the transaction list 308 by setting the relevantdata field to null. In state 684, the participant module frees the spaceallocated to the relevant transaction descriptor 304. In state 686, theparticipant module frees the space allocated to the relevant blockdescriptor 306. In state 688, the participant module frees the spaceallocated to the relevant journal block 190.

VII. Exemplary Replay

When necessary, a nonlinear journal, such as nonlinear journal 300, maybe replayed to the respective hard-disk drives 150. In other words, thecontents of the journal blocks 190 may be written to the correspondingdisk blocks 178. In the illustrated embodiment, journal blocks 190corresponding to committed and done transactions are replayed, andjournal blocks 190 corresponding to prepared transactions are ignored.

To support down drives, in one embodiment, mount (rebuild) and replayare separated into two stages. Down drives are discussed in more detailbelow. On initial mount, the nonlinear journal 300 is traversed startingfrom the journal descriptor 302 (super-block). All transactiondescriptors 304 and block descriptors 306 are identified, and theirinformation is remembered. With the information obtained from traversingthe nonlinear journal 300, the global transactions are resurrected insystem memory 152. After waiting to resolve them, the hard-disk drivesare replayed on a per-drive basis. Initially, the hard-disk drives 150are effectively down. A set of active (not dead) hard-disk drives 150 isthen provided, and any memory blocks 180 not associated with an activedrive are discarded. Once the journal is mounted, replay may be calledon each drive as it is brought from down to up. FIGS. 7A and 7B describerebuild and replay in greater detail.

FIG. 7A illustrates a flowchart of one embodiment of rebuildingtransactions from nonlinear journal 300. In state 702, the participantmodule traverses a nonlinear journal 300 starting at the journaldescriptor 302 (super-block). As the participant module traverses thenonlinear journal 300, it remembers information about the transactiondescriptors 304 and block descriptors 306. In state 704, the participantmodule resurrects the global transactions based on the gatheredinformation from the nonlinear journal 300. In one embodiment, theglobal transactions are reconstructed in system memory 152. The globaltransactions are restored to their respective states, as indicated bythe txn_state data field 356 of their corresponding transactiondescriptor 304. In state 706, the participant module waits until thetransactions resolve. In state 708, the participant module replays thehard-disk drives 150 on a per-drive basis, which is illustrated ingreater detail in FIG. 7B.

FIG. 7B illustrates a flowchart of one embodiment of replaying, on aper-drive basis, transactions resurrected from nonlinear journal 300, torespective hard-disk drives 150. In state 752, the participant moduledetermines whether a particular hard-disk drive 150 is available forstorage. In one embodiment, the participant module may determine that aparticular hard-disk drive 150 is available for storage upon receiving acommunication from the respective hard-disk drive 150. In anotherembodiment, the participant module may actively query whether ahard-disk drive 150 is available (up) or unavailable (down) for storage.If the relevant hard-disk drive 150 is unavailable for storage (down),then the participant module proceeds to the end. If the relevanthard-disk drive 150 is available for storage (up), the participantmodule repeats states between 754 and 760 for the transactions blockscorresponding to the relevant hard-disk drive 150. In state 758, theparticipant module replays the relevant blocks of the relevanttransaction to the relevant hard-disk drive 150. In the illustratedembodiment, the in-memory buffers, such as memory blocks 180, arewritten (flushed) to the corresponding hard-disk drives 150. In otherembodiments, the hard-disk drives 150 may be replayed directly with thecontents of corresponding journal blocks 190.

VIII. Exemplary Down Drive Support

A hard-disk drive, such as one of hard-disk drives 150, may go out ofservice (or down) upon a temporary failure, including, but not limitedto, a timeout or a cabling error. In one embodiment, hard-disk drives,such as hard-disk drives 150, may go down, and a journal, such asnonlinear journal 300, may retain journal blocks 190 corresponding tothe down drive, and the drive may be safely brought back into servicewhile a distributed system, such as distributed system 100, is stilloperational.

FIG. 4E illustrates a state diagram describing one embodiment ofproviding support for down drives in nonlinear journal 300. In oneembodiment, hard-disk drives, such as hard-disk drives 150, may not beread or written in the down state. Moreover, upon transition to the downstate, transactions in the writing state (R) are aborted; whereas,transactions in the prepared state (P), committed state (C), and donestate (D) are retained. Committed transactions may not flush blocks todown drives, but remain instead in the journal, such as nonlinearjournal 300. Transactions may behave normally in the presence of downdrives. Prepared transactions may transition to committed on a commitcommand and to unknown on an abort command. Committed transactions maystill transition to done on a done command. Transactions stay in thedone state because a down drive does not flush its blocks. When a downdrive transitions to up, the journal blocks, such as journal blocks 190,associated with committed or done transactions are replayed to the drivebefore it changes state. When a drive transitions to dead, it may notreturn. Transactions in the writing state (W) are aborted, and anyjournal blocks, such as journal blocks 190, associated with the driveare freed.

The following are exemplary entry points into one embodiment of ajournal subsystem for bringing drives up and down:

-   -   drive_up( ): Bring a drive up from the down state.    -   drive_down( ): Bring a drive down from the up state.    -   drive_dead( ): Move a drive to the dead state from either the up        or down states.

The following exemplary pseudocode further describes one embodiment of ajournal subsystem providing support for bringing drives up and down:

Function kill( ):  for each b in Dp U Dp:   block_abort(b)  DF = DP = ∅ set state to dead Function abort_writing( ):  for each txn in state W:  abort (txn) Function replay(txn):  Write all blocks associated withthis drive on this transaction to the disk. in state up:  on drive_down(): abort_writing( ) set state to down  on drive_dead( ): abort _writing() kill( ) in state down:  on drive_up( ): for each txn in state CUD: replay (txn) set state to up  on drive_dead( ): kill( ) In state dead:

IX. Exemplary Support for Shadow Buffers

FIG. 8 illustrates embodiments of data structures implementing anonlinear journal capable of maintaining a “shadow buffer.” When amemory block 180 is overwritten by a new transaction before the memoryblock 180 has been written (flushed) to the corresponding hard-diskdrive 150, there is a need to keep a copy of the value of the memoryblock 180 in the event that the overwriting transaction aborts. In oneembodiment, a copy of the overwritten value is stored in system memory152. In another embodiment, the journal block 190 corresponding to theoverwritten value is preserved in the nonlinear journal 300 as a “shadowbuffer.” If the overwriting transaction commits, then the journal block190 serving as the shadow buffer is freed. If the overwritingtransaction aborts, however, the memory block 180 is restored to thevalue saved in the journal block 190 (the shadow buffer).

In the illustrated embodiment, two transactions journaled in persistentstorage 156 modify the same storage destination. Transaction T_(A) andtransaction T_(B) both include journal blocks 190 that reference thesame disk block [1] 178 on hard-disk drive [1] 150. If the firsttransaction, transaction T_(A), is not written (flushed) before thesecond transaction, transaction T_(B), this circumstance gives rise tothe need to create a shadow buffer for the previous data in therespective memory block 180. Nonlinear journal 300 illustrates oneembodiment of a nonlinear journal that may be implemented to provide ashadow buffer. This embodiment reduces the need to save a copy of thedata in system memory 152, thereby reducing the time and space expenseof keeping a system-memory copy.

The following exemplary pseudocode further describes one embodiment ofkeeping a shadow buffer in the nonlinear journal 300:

write_buf( ) if (b is dirty):  grab a reference on the previous transaction's block modify contents of b write b to journal commit( )delete previous transaction's block out of the journal mark b dirtyabort( ) restore previous contents of b by reading data from previoustransaction's block mark b dirty

A. Example Shadow Buffers

FIGS. 9A, 9B and 9C illustrate embodiments of maintaining a shadowbuffer in a nonlinear journal. In 900, the memory block 180corresponding to disk block [7] 178 on hard-disk drive [1] 150 is“Invalid,” meaning that there is no memory block 180 that correspondscurrently to disk block [7] 178 on hard-disk drive [1] 150. Thenonlinear journal 300 also has no journal block 190 that corresponds todisk block [7] 178 on hard-disk drive [1] 150. Disk block [7] 178 onhard-disk drive [1] 150 stores the value “XYZ.”

In 902, the participant module reads the value of disk block [7] 178 onhard-disk drive [1] 150 into a memory block 180 corresponding to diskblock [7] 178 on hard-disk drive [1] 150. The system memory 152 now hasa valid memory block 180 corresponding to disk block [7] 178 onhard-disk drive [1] 150.

In 904, the participant module receives a “write [abc] to drive 1,disk_block 7” message corresponding to transaction T_(A). Theparticipant module stores the value “abc” in the memory block 180corresponding to disk block [7] 178 on hard-disk drive [1] 150. Theparticipant module then stores the data “abc” to a journal block 190.The respective memory block 180 is clean, which means its data contentscannot yet be written to hard-disk drive [1] 150.

In the illustrated embodiment, when a memory block 180 is “clean,” thejournal subsystem forbids the disk subsystem from writing the contentsof the memory blocks 180 to their corresponding hard-disk drives 150.When a memory block 180 is “dirty,” the journal subsystem permits thedisk subsystem to write the contents of the memory blocks 180 to theircorresponding hard-disk drives 150. In the illustrated embodiment,memory blocks 180 are clean when their contents are uncommitted. Inother words, when a memory block 180 is first written with data from anuncommitted transaction, the memory block 180 is clean. After thecorresponding transaction is committed, the memory block 180 becomesdirty, and its contents may be written to the respective hard-disk drive150. If a memory block 180 is written before its correspondingtransaction is committed, the result is data corruption because thetransaction could abort, leaving the contents of an aborted transactionstored on the respective hard-disk drive 150. After a memory block 180becomes dirty, it becomes clean again when its contents are written(flushed) to the respective hard-disk drive 180, or when the contents ofa new transaction are written to the memory block 180. In FIGS. 9A, 9B,and 9C, the contents of memory block 180 corresponding to disk block [7]178 on hard-disk drive [1] 150 are marked with a subscript “d” toindicate when the memory block 180 is dirty. When memory block 180 isclean, its contents are not marked.

In 906, the participant module receives a “commit T_(A)” message. Theparticipant module atomically writes the txn_state data field 356 oftransaction descriptor [T_(A)] 304 in the nonlinear journal 300 to“committed.” Afterwards, the respective memory block 180 is now dirty,indicating that the memory block 180 may be written to hard-disk drive[1] 150. The respective disk block 178 still retains the value “XYZ.”

In 908, the participant module writes (flushes) the value “abc” tohard-disk drive [1] 150. The effect is that the drive cache [1] 172includes a cache block [7] 176 with the value of “abc.” Whether the diskcontroller 170 has written the data “abc” to the corresponding diskblock [7] 178 on the appropriate disk platter 174 is uncertain. Afterflushing the contents of the respective memory block 180 to hard-diskdrive [1] 150, the respective memory block 180 is now clean, and therespective disk block [7] 178 on hard-disk drive [1] 150 stores eitherthe value “abc” or “XYZ.”

In state 910, the participant module syncs hard-disk drive [1] 150,meaning that the contents of drive cache [1] 172 are flushed tohard-disk drive [1] 150. Accordingly, disk block [7] 178 of hard-diskdrive [1] 150 is now synchronized with the respective journal block 190,and the respective journal block 190 is now freed from the nonlinearjournal 300. The respective memory block 180 is still clean, and therespective disk block 178 now stores the data value “abc.”

Although in the illustrated embodiment synchronization occurs on aper-drive basis, in other embodiments synchronization may transpire on aper-block, per-cache, or per-platter basis, as well as other suitablealternatives and equivalents. For example, an embodiment that includedhard-disk drives that offer “forced unit access” could be synchronizedon a per-block basis. Additionally and/or alternatively, an embodimentthat included hard-disk drives comprising separate caches for a bundleof disk platters could be synchronized on a per-cache or per-platterbasis. Moreover, although in the illustrated embodiment synchronizationis the result of an explicit cache flush, in other embodimentssynchronization may include, for example, periodically reading the diskplatters 174 to confirm that the relevant disk blocks 178 have beenwritten. The corresponding hard-disk drives 150 would be configured toreturn the contents of its respective disk blocks 178, rather than theassociated cache blocks 180, which may not have been flushed yet.

In still other embodiments, synchronization may be based on a waitingperiod that guarantees cache flushing within the determined time frame.For example, hard-disk drives 150 may guarantee a cache flush on aperiodic basis. A participant module could be configured to notify thejournal subsystem of the cache flush after receiving notification fromthe respective hard-disk drive 150 of a cache flush or as part of acache-flush schedule that the hard-disk drives 150 follow on adeterminative basis. In yet other embodiments, a least recently used(LRU) expiration model may be used in addition to, or in place of, thesynchronization expiration model.

FIG. 9B illustrates an example use of a shadow buffer. Items 900, 902,904, and 906 in FIG. 9B are the same as in FIG. 9A. In 908, therespective memory block 180 corresponding to disk block [7] 178 onhard-disk drive [1] 150 is not, however, written (flushed) to hard-diskdrive [1] 150. Instead, in 912, a new transaction T_(B) overwrites therespective memory block 180 and also writes a corresponding journalblock 190. Because a memory block 180 was overwritten before theprevious value was written to hard-disk drive [1] 150, the participantmodule keeps a copy of the previous value.

In one embodiment, nonlinear journal 300 keeps the copy of the previousvalue. Because transaction T_(A) has committed, its associated journalblocks 190, including the journal block 190 corresponding to theprevious value of the memory block 180, are candidates for being freedfrom the nonlinear journal 300 after a drive sync for a correspondinghard-disk drive 150, such as hard-disk drive [1] 150. To keep a shadowbuffer of the overwritten memory block 180, the participant module keepsa reference to the journal block 190 corresponding to the value of theoverwritten memory block 180. The associated block descriptor 306 andtransaction descriptor 304 are likewise retained in the journal for aslong as needed. The participant module keeps the shadow buffer until theoverwriting transaction T_(B) either commits or aborts. If thetransaction T_(B) commits, there is no need to keep the shadow bufferbecause the respective memory block 180 will not be rolled back to theprevious value, and the shadow buffer and the associated descriptors arecandidates for being freed. If the transaction T_(B) aborts, therespective memory block 180 is rolled back to the previous value, usingthe shadow buffer in the nonlinear journal 300, and the journal block190 that served as the shadow buffer becomes the journal block 190 forthe rolled-back value of the memory block 180.

In 912, the respective memory block 180 is clean, the nonlinear journal300 retains the journal block 190 with the previous data value “abc,”the nonlinear journal 300 also retains the journal block 190 with thenew value “hello,” and disk block [7] 178 on hard-disk drive [1] storesthe value “XYZ.”

In 914, the transaction T_(B) commits. This causes the respective memoryblock 180 to become dirty, meaning its contents may be written tohard-disk drive [1] 150. The shadow buffer in the nonlinear journal 300does not need to be retained, and the respective journal block 190 withthe value “abc” can be freed in a similar manner as if a drive sync hadoccurred. In the illustrated embodiment, the journal block 190 thatserved as the shadow buffer is freed automatically. In otherembodiments, the journal block 190 that served as the shadow buffer mayremain in persistent memory 152 until a garbage collector collects it.The nonlinear journal 300 retains, however, the journal block 190 withthe overwritten data value “hello.” Disk block [7] 178 on hard-diskdrive [1] 150 still stores the value “XYZ.”

In 916, the participant module writes (flushes) the respective memoryblock 180 to hard-disk drive [1] 150. The respective memory block 180 isnow clean, and disk block [7] 178 on hard-disk drive [1] 150 is in anunknown condition, storing either the value “XYZ” or the value “hello,”depending on whether or not the drive cache [1] 172 has flushed itscontents to disk platter [1] 174.

In 918, the participant module syncs hard-disk drive [1] 150. At thispoint, the memory block 180 is still clean, the journal block 190 withthe value “hello” is now freed, and disk block [7] 178 on hard-diskdrive [1] 150 stores the value “hello.”

FIG. 9C illustrates another alternative, showing the effects of an“abort” message after a shadow buffer has been created. Both 900, 902,904, and 906 illustrated in FIG. 9C are the same as those illustrated inFIGS. 9A and 9B. And 912 is the same as 912 illustrated in FIG. 9B.Following 912, the respective memory block 180 stores the new data value“hello,” and the data has not yet been written (flushed) to hard-diskdrive [1] 150, making the memory block 180 clean. The nonlinear journal300 includes a journal block 190 with the previous value “abc,” which isthe shadow buffer. The nonlinear journal 300 also includes a journalblock 190 with the new value “hello,” and disk block [7] 178 onhard-disk drive [1] 150 stores the value “XYZ.”

In state 920, the participant module receives an “abort (T_(B))”message. Because the overwriting transaction has aborted, the respectivememory block 180 is rolled back to the previous value “abc” using theshadow buffer in the nonlinear journal 300. The journal block 190 thathad stored the aborted value “hello” may be freed. In the illustratedembodiment, the journal block 190 that kept the shadow buffer isautomatically freed because aborted journal blocks 190 are treated as ifthey never happened. In other embodiments, the aborted journal blocks190 may be freed by a garbage collector. Disk block [7] 178 on hard-diskdrive [1] 150 stores the value “XYZ.” The memory block 180 now has thedata value “abc.” Because the corresponding transaction (T_(A)) hasalready committed and because the restored contents of the memory block180 have not been written to hard-disk drive [1] 150, the memory block180 is dirty.

In 922, the participant module writes (flushes) the respective memoryblock 180 to disk block [7] 178 on hard-disk drive [1] 150. Therespective memory block 180 still has the data value “abc,” but is nowclean. The nonlinear journal 300 still retains the journal block 190with the value “abc.” Disk block [7] 178 on hard-disk drive [1] 150stores the value of either “abc” or “XYZ,” depending on whether thedrive cache [1] 172 has flushed its contents to disk platter [1] 174.

In 924, the participant module syncs hard-disk drive [1] 150. Therespective memory block 180 still stores the value “abc” and is stillclean. The respective journal block 190 with the data value “abc” isfreed, and disk block [7] 178 on hard-disk drive [1] 150 stores thevalue “abc.”

B. Example Shadow Buffer Procedure

FIG. 10 illustrates one embodiment keeping a shadow buffer in nonlinearjournal 300. In state 1002, the participant module determines whetherthe relevant memory block 180 is dirty. For purposes of this example,the relevant memory block 180 is a memory block 180 that overwritesprevious data with new data. If the relevant memory block 180 is dirty,then the participant module proceeds to state 1005. If the memory block180 is clean, or not dirty, then the participant module proceeds tostate 1004. In this example, a clean memory block 180 refers to a memoryblock 180 whose contents have committed and have already been written todisk. Memory blocks 180 whose contents are uncommitted are not relevantto shadow buffers because the need for a shadow buffer arises when amemory block 180 is overwritten, and a memory block with uncommittedcontents may not, in the illustrated embodiment, be overwritten.

In state 1004, the participant module keeps a reference to the journalblock 190 whose contents correspond to the previous data of the relevantmemory block 180. This journal block 190 is referred to as a “shadowbuffer” for the relevant memory block 180. In state 1006, theparticipant module modifies the relevant memory block 180 with the new,overwriting data. In state 1008, the participant module writes the newdata as a new journal block 190. In state 1010, the participant moduledetermines whether the overwriting transaction either commits or aborts.In one embodiment, the participant module determines whether thetransaction commits or aborts by waiting to receive a “commit” or“abort” message from, for example, a global transaction module. If theoverwriting transaction aborts, then the participant module proceeds tostate 1014. If the overwriting transaction commits, then the participantmodule proceeds to state 1016. In state 1014, the participant modulerestores the relevant memory block 180 to the previous data using theshadow buffer, the journal block 190 with the previous data contents ofthe relevant memory block 180. In state 1016, the participant moduleremoves the shadow buffer, which is the journal block 190 correspondingto the previous data.

If, in state 1002, it is determined that the relevant memory block 180is clean, then the participant module, in one embodiment, does not needto keep a shadow buffer. The procedure for processing an overwritewithout a shadow buffer is similar to the procedure described above,except the participant module need not keep track of a shadow buffer. Instate 1005, the participant module overwrites the relevant memory block180 with the new (overwriting) data. In state 1007, the participantmodule writes the new data to another journal block 190 beside thejournal block 190 that keeps the previous data. The journal block 190with the previous data is now a candidate for being released from thenonlinear journal 300. If the overwriting transaction aborts before thejournal block 190 with the previous data is synced, the participantmodule can use its contents to restore the relevant memory block 180. Ifthe overwriting transaction aborts after the journal block 190 with theprevious data has been synced, the participant module can read thecontents of the corresponding disk block 178 to restore the relevantmemory block 180. In state 1009, the participant module determineswhether the overwriting transaction commits or aborts. In oneembodiment, the participant module may make the determination by waitingto receive either a “commit” or “abort” message from, for example, aglobal transaction module. If the overwriting transaction commits, theparticipant module proceeds to the end. If the overwriting transactionaborts, then the participant module proceeds to state 1015. In state1015, the participant module disregards the contents of the relevantmemory block 180, which contents correspond to the aborted transaction.

X. Supporting Concurrent Transactions with a Nonlinear Journal

In one embodiment, a nonlinear journal, such as nonlinear journal 300,allows for limited concurrent transaction support. In some embodiments,when the participant module attempts to write transactions that bothinclude a journal block 190 corresponding to the same disk block 178 onthe same hard-disk drive 150, each transaction owns the blockexclusively until it resolves. Thus, the first transaction acquires theblock and holds it from the moment it writes it until the transaction iscommitted. At this point, another transaction may wake up and acquirethis block for writing. In other words, in some embodiments, twotransactions may not both be in the prepared state with data writes tothe same block.

To allow for a limited concurrency, some embodiments of a nonlinearjournal, such as nonlinear journal 300, take advantage of the idea thatsome operations are commit order independent. For example, a sequence ofadditions and subtractions will yield the same result regardless of theorder of the operations. Thus, as long as the nonlinear journal 300includes enough data to reconstruct the block, there is no requirementfor exclusivity for these operations. The data structures and modulessupported by the nonlinear journal 300 to implement concurrentoperations are referred to herein as Deltas. In one embodiment, Deltascan be used at the same time as block-level writes, but since blockwrites are order-dependent, block writes may wait for delta transactionson the block to commit and flush, and new delta writes may wait for theblock write.

There are a number of operations that can be supported using Deltas.Primarily, these operations are order independent, but Deltas may alsobe used for some partially ordered operations. In one embodiment, orderindependent operations may include: addition, subtraction, integermultiplication, maximum, minimum, XOR, set union and set intersection.In one embodiment, some of these operations, such asaddition/subtraction, may be strictly reversible, meaning that at anypoint a transaction may be removed by inverting the operation andreapplying it. In one embodiment, other operations (such as maximum) arenot strictly reversible, and reapply the delta operations from someknown point. In one embodiment, some operations (such as integermultiplication and set union) are subject to overflow. If an overflowoccurs, the result may be undefined. In one embodiment, it is theresponsibility of the user of the nonlinear journal 300 to guaranteethis does not occur. In one embodiment, a type of delta operation may beincompatible with other types of delta operations, as well as with blockoperations. In one embodiment, the caller guarantees that the operationsdo not overlap.

Another type of operation that might be done with delta operations is ofthe class of partially ordered block writes. For example, consider a512-byte block which is split into four 128-byte regions. If apredecessor block exists, a delta operation might represent apartial-block write to any of these four sub-blocks, each of which wouldbehave similarly to a normal write to a whole block. This allowssubdivision of a block between multiple users. Each of these sub-blockscould be written independently of the other sub-blocks. Writes to thissub-block would therefore be partially-ordered such that the ordering isguaranteed only within sub-block boundaries. Up to four transactionscould be in progress on a four-region block of this sort at the sametime without contention. The “delta” operation in the case of apartially-ordered operation would simply be to overwrite the sub-blockin question.

The idea of partial-block overwrites being independent could be extendedto include order-independent operations as well. The basic principle isthat a type of delta operation is incompatible with another at the samelocation. Thus, a single block might be subject to partial-overwritesand delta operations at the same time provided that the same operationsdo not overlap. One way this could be implemented would be with a maskoverwrite. Instead of just overwriting a simple region, a partial-blockwrite would overwrite the block through some arbitrary mask, which masksout the portions of the block meant to be modified through other deltaoperations. In this way, a large number of transactions couldefficiently apply delta operations concurrently, while a disjoint set oftransactions could mask overwrite the block serially. To do a full blockoverwrite, a lock that excludes both mask and delta writes may have tobe taken. FIG. 15 illustrate one embodiment of combining commit orderindependent operations with partially ordered block writes.

In one embodiment, to avoid having to support 10 errors on reads, Deltasmaintain the property that journal replay is write-only to the hard-diskdrives 150. Because of this, delta operations stored in the nonlinearjournal 300 points back to a copy of a journal block 190—not a diskblock 178. Thus, in one embodiment, the journal does not apply a Deltaunless it can find a predecessor journal block 190. When applying aDelta, the nonlinear journal 300 attempts to find a predecessor journalblock 190 for the Delta. If one exists, the Delta will be applied. If nojournal block 190 corresponding to the same disk block 178 as the Deltaexists in persistent memory 152, the journal module will read the diskblock 178 and then apply the Delta in a corresponding memory block 180,writing the modified memory block 180 as a normal block-write, whichincludes recording a journal block 190. This allows the next deltaoperation to find a predecessor block—the respective journal block 190.

In another embodiment, the predecessor journal block 190 may be omitted,and Deltas may be applied to the data value at the corresponding diskblock 178, which requires reading the contents of the disk block 178into a memory block 180, performing the operation with the Delta, andthen writing out the result back to the disk block 178. In addition tothe general expense of 10 operations, the read may introduce thepossibility of a read error, which would prevent the Delta from beingwritten to the disk block 178. In some embodiments, this problem may beovercome by treating the read error as a permanent error and discardingthe remaining Deltas until the disk block 178 is overwritten. Theseembodiments may require modifications to other embodiments describedherein.

There may be many important applications of Deltas to supportcommit-order independent and partially ordered operations. Some possibleuses include, without limitation, updating data that representsaccounting, ctime, parity, combinations of the same, and suitableequivalents.

A. Delta Data Structures

FIGS. 11A and 11B illustrate embodiments of data structures forimplementing a nonlinear journal capable of handling concurrenttransactions. In the illustrated embodiment, nonlinear journal 300includes at least three transaction descriptors 304, corresponding tothree separate transactions that have been written to nonlinear journal300. The nonlinear journal 300 also includes a block descriptor 306 thatincludes at least one journal block 190. The nonlinear journal 300,illustrated in FIG. 11A, also includes at least two delta descriptors1102. Delta descriptors 1102 include metadata representing certainconcurrent transactions. In the illustrated embodiment, the two deltadescriptors 1102 and the block descriptor 306 correspond to the samedisk block [27] 178 on hard-disk drive [4] 150.

In the illustrated embodiment delta descriptors 1102 are chained alongwith the block descriptors 306 of each transaction descriptor 304. Adelta descriptor 1102 includes a desc_link data field 366 for linking toit other delta descriptors 1102 or block descriptors in a block list310. A delta descriptor also includes multiple groups of five datafields (called delta elements) that include: a drive data field 368, adisk_block data field 370, an off-set data field 1104, an operation datafield 1106, and a value data field 1108. In the illustrated embodiment,there are no direct links between delta descriptors 1102 correspondingto the same disk block 178 on the same hard-disk drive 150, or betweendelta descriptors 1102 and block descriptors 306. Linkage of thesedescriptors is determined implicitly when the journal is mounted basedon the drive and block addresses. In general, if a Delta exists in thenonlinear journal 300, some journal block 190 to the same locationexists previously for the Delta to apply to it. As described above withreference to FIG. 7A, during the mount (rebuild) process, the nonlinearjournal 300 is scanned, global transactions are resurrected based on thejournal blocks 190 and the associated descriptors, the system waitsuntil the transactions have resolved, and the hard-disk drives 150 arethen replayed.

Because, in the illustrated embodiment, a Delta cannot exist by itself,requiring a predecessor block to chain off of; expiring Deltas from thenonlinear journal 300 may require an atomic operation. For the Deltas tostop being relevant and be safe to delete, some atomic operation is donewhich reliably makes them irrelevant. If the journal block 190 that therespective Delta is associated with is written to disk and synced(presumably after a period of inactivity), then, in one embodiment, thepredecessor block may be unlinked, provided that the previouspredecessors have already been unlinked. At this point, the Deltas haveno block to refer to and can be unlinked as well. On replay, they cansimply be ignored.

If the delta block is written to often, though, it may never be writtento disk. In this case, it may not be safe to delete the predecessor, soa method to make it safe to delete the deltas and/or predecessor may beused. In one embodiment, the delta writers are quiesced. The delta blockcould be exclusively locked and synchronously written to the respectivehard-disk drive 150. Unfortunately, this may require blocking deltas fora long period of time—which may block the writers in the cluster forsome applications. Embodiments of systems and methods for quiescing aredisclosed in U.S. patent application Ser. No. 11/357,740, titled“Systems and Methods for Providing a Quiescing Protocol,” filed on Feb.17, 2006, which is hereby incorporated by reference herein in itsentirety.

In another embodiment, not involving disk IO, a sequence of Deltas couldbe collected together and written out in a new Delta against the sameblock. If, for example, a predecessor journal block 190 had twocommitted Deltas among a chain of five total Deltas, the two committedDeltas could be collected together and written as a combined Deltaalready committed when it is written. It may be difficult, however, toatomically unlink the two old Deltas. Thus, this new “compacted” deltatransaction may need to include some sort of blackout references in it,indicating that the previous delta transactions no longer apply. If theparticipant module later reads the delta chain, the blackout referencessignal it to ignore those Deltas, causing the participant module to readthe combined Delta, but not the two lingering Deltas that have beenblacked out, though not atomically unlinked.

In still another embodiment, the journal module may do a block-wiseoverwrite of the entire journal block 190 that the Delta applies to.Once a new definitive copy of the journal block 190 exists, the previouspredecessor journal block 190 and delta chains can be freed from thejournal safely without worry of atomicity. This new journal block 190may be thought of as a pseudo-transaction, combining the operations ofmultiple transactions that apply to the same disk block 178.

Unlinking a delta from the nonlinear journal 300 once it is masked by afull block write (and thus safe), in some embodiments, may approximateunlinking a journal block 190. In some embodiments, an element of thedelta descriptor 1102 is atomically zeroed, unless doing so would leavethe entire block descriptor 306 and/or the transaction descriptor 304empty, in which case unlink writes may be elided.

In some embodiments, some blocks may be written entirely through Deltas,and never (or almost never) through whole block writes. Since a wholeblock write is necessary to free a delta from the journal, some form ofperiodic flush, in some embodiments, may be necessary or the journal mayeventually become completely full of delta blocks. To perform a periodicflush efficiently, the journal may construct a definitive copy of theblock to write out as part of a dummy transaction. This block includes acopy of the block as it would appear after all committed transactions,but without the changes that may be made by writing transactions, whichmay abort. For strictly reversible operations, this block may beconstructed by taking the current copy of the block, and then reversingall writing deltas and un-applying them. This would yield the lastdefinitive copy. However, this technique, in some embodiments, may onlybe possible for reversible operations.

A general technique which may work for all order-independent andpartially-ordered operations is to restore the last predecessor blockinto a temporary buffer, and then apply each committed delta to thisblock. This may be much more expensive than the strictly reversibleapproach in terms of CPU time. Once the definitive copy block isconstructed, it is written to the journal and then linked into theproper place. Because it represents an order-dependent operation on theblock, it is linked into the transaction list after the last transactionit masks, but before any transactions including deltas that it does notrepresent. In one embodiment, the participant module simply waits forall transactions in the delta chain to either commit or abort, andblocks other transactions from starting, until the definitive copy isconstructed and written. In another embodiment, the transactiondescriptor 304 could be inserted into the transaction list 308 after thelast committed transaction via two ordered atomic writes. This“splicing” of the definitive copy results in this pseudo-transactioncommitting out of order. Once the dummy transaction is committed, insome embodiments, all blocks and deltas which it masks may be freed fromthe journal, as if they had been overwritten. For example, if apredecessor journal block 190 had five Deltas, two of which werecommitted, a definitive copy could be spliced just after the twocommitted Deltas, but before the three others. The definitive copybecomes the new predecessor journal block 190 and the two committedDeltas can be released as if they had been overwritten.

It should be noted that journal deltas as described may requireallocating space in the journal as part of the process of freeing spacein the journal. Thus, in some embodiments, there may be a method where,when a journal delta is created, enough space in the journal is reservedto represent a definitive block which allows the deltas and predecessorblocks to be freed. If a drive is down, compaction via definitive blockcan still be applied, since it only requires journal IO, not disk IO.Thus, once a drive is down, in some embodiments, any block and deltasequence can be compacted into a single block in the journal.

B. Example Concurrent Transactions

FIGS. 12A, 12B, 12C and 12D illustrate embodiments implementingconcurrent transactions in a nonlinear journal. In 1200, the relevantmemory block 180 is in an “Invalid” state with respect to disk block[27] 178 on hard-disk drive [4] 150. In other words, the system memory152 does not have a memory block 180 with the current value of therelevant disk block. Disk block [27] 178 on hard-disk drive [4] 150stores the value of “1.” In 1202, the participant module executes arequest for an order-independent operation, transaction T₁, on thecurrent value of the disk block [27] 178 on hard-drive disk [4] 150.Because the memory block 180 is in the “Invalid” state, the participantmodule first reads the value of disk block [27] 178 on hard-disk drive[4] 150. Following the read, the relevant memory block 180 stores thevalue “1,” and disk block [27] 178 on hard-disk drive [4] 150 alsostores the value “1.” The participant module then adds the value “9” tothe value in the relevant memory block 180, yielding a result of 10,which is stored in the relevant memory block 180. Because there is nopredecessor journal block 190 for disk block [27] 178 on hard-disk drive[4] 150, the participant module then writes the value of the data storedin memory block 180 to a journal block 190, storing as a result thevalue “10.” Because transaction to which journal block 190 correspondshas not yet committed, journal block 190 is writing, which is denotedwith a subscript “w.” In FIGS. 12A, 12B, 12C, and 12D, the subscript “w”on a journal block 190 represents the writing state and the subscript“c” represents the committed state.

In 1204, the participant module receives a “commit (T₁)” message from aglobal transaction module. The state of the relevant memory block 180becomes “dirty.” In state 1206, the participant module receives a secondwrite request, transaction T₂, from a global transaction module. Theparticipant module adds the value “2” to the respective memory block180, yielding a result of 12, which is then stored in memory block 180.The participant module then writes to persistent storage 156 a deltaelement representing the order-independent operation of adding the value“2.” Because there is a predecessor block, the journal block 190corresponding to disk block [27] 178 on hard-disk drive [4] 150, theparticipant module may write a delta element for transaction T₂.

In 1208, the participant module receives a request for a thirdorder-independent operation, transaction T₃, on disk block [27] 178 onhard-disk drive [4] 150. The participant module adds the value “5” tothe respective memory block 180, and writes an appropriate delta elementto persistent memory 156.

In 1210, the participant module receives a request for a fourthorder-independent operation, transaction T₄, on disk block [27] 178 onhard-disk drive [4] 150. The participant module adds “3” to therespective memory block 180, storing the value “20” in the respectivememory block 180. The participant module also writes an additional deltaelement to persistent memory 156.

In 1212, a participant module receives an “abort (T₂)” message. In oneembodiment, an abort procedure, the participant module reads the oldpredecessor journal block 190 from persistent memory 156 and, for eachdelta that has not aborted, applies the delta, using a copy of the deltakept in system memory 152. In another embodiment, for strictlyreversible operations, the operation is reversed and the reversed deltais applied. In still another embodiment, the deltas are read frompersistent memory 156, rather than from separate copies in system memory152. The participant module subtracts “2” from the respective memoryblock 180, and frees the delta element from transaction T₂. In oneembodiment of the abort procedure, the transaction descriptor 304 isunlinked and then the associated blocks, deltas, and descriptors arereverted in the persistent memory 156 and the space is freed.

In 1214, the participant module receives a “commit (T₃)” message.Because not all of the delta elements have been committed for disk block[27] 178 on hard-disk drive [4] 150, the memory block 180 remains in anundirty state. In state 1216, the participant module receives a “commit(T₄)” message. Because the predecessor journal block 190 and thecorresponding delta elements are committed, the respective memory block180 enters the dirty state.

FIG. 11B illustrates the status of the nonlinear journal 300, withrespect to disk block [27] 178 on hard-disk drive [4] 150, followingstate 1216. Three transaction descriptors [T₁, T₃, T₄] 304 are linkedinto the transaction list 308. Transaction descriptor [T₁] 178 links toblock descriptor [B₁] 306, which links to the predecessor journal block190 for disk block [27] 178 on hard-disk drive [4] 150. Transactiondescriptors [T₃, T₄] 178, respectively, link to delta descriptors [D₃,D₄] 1102, which include delta elements corresponding to disk block [27]178 on hard-disk drive [4] 150.

FIGS. 12A, 12B, 12C, and 12D illustrate the states until this point.FIG. 12A further illustrates a drive sync with deltas. In state 1217,the participant module writes (flushes) the contents of memory block 180to hard-disk drive [4] 150. The contents of memory block 180 are nowwritten into a respective cache block 176 of drive cache 172 onhard-disk drive [4] 150. After writing its contents, memory block 180becomes clean. In 1218, the participant module syncs hard-disk drive [4]150. The disk block [27] 178 on hard-disk drive [4] 150 now has the datavalue of “18.” At this point, the journal block 190 and thecorresponding delta elements may be unlinked from the nonlinear journal300 and their space freed in persistent memory 156. In one embodiment, adelta element may be freed directly. A portion of the delta element isatomically overwritten to indicate that it is free, or, if it is thelast delta in the delta descriptor 1102, the entire delta descriptor1102 is freed.

In 1220, the participant module receives a message from the globaltransaction module requesting another transaction T₅, which alsoincludes the respective memory block 180. The value 10 is added to therespective memory block 180 and the new value of the memory block 180“28” is written to the nonlinear journal 300. In 1222, the participantmodule receives a message from the global transaction module committingtransaction 5. The respective memory block 180 is in the dirty state,the journal block 190 is unchanged, and this block 178 is alsounchanged.

FIG. 12B further illustrates coalescing predecessor blocks and deltaelements. In 1224, the participant module receives a fifth request foran order-independent operation, transaction T₅, on disk block [27] 178on drive [4] 150. The participant module adds the value “10” to therelevant memory block 180, and writes a corresponding delta element topersistent memory 156. In state 1226, the participant module receives a“commit (T₅)” message. The relevant memory block 180 is now in a dirtystate. In 1228, the participant-module receives a message to overwrite,transaction T₆, the relevant memory block 180. This may be anindependent overwrite request, but it may also be a system-generatedoverwrite to coalesce the previous delta elements into a new predecessorjournal block 190. The write value “28” equals the combination of theprevious predecessor journal block 190 and its associated deltaelements.

Because the previous value of the respective memory block 180 was notyet written to hard-disk drive [4] 150, the participant module creates ashadow buffer in the nonlinear journal 300 in order to save a copy ofthe previous value in the event that transaction T₅ aborts. In 1230, theparticipant module receives a “commit (T₆)” message. The relevant memoryblock 180 is in a dirty state, and there is no need to retain a copy ofthe previous value of the memory block 180, so the shadow buffer in thenonlinear journal 300 is freed.

FIG. 12C further illustrates overwriting a predecessor block and itsassociated delta elements. In 1232, the participant module receives amessage to overwrite the relevant memory block 180 with the value “17.”The participant module overwrites the respective memory block 180.Because the previous value of the respective memory block 180 “18” wasnot written to disk yet, the participant module creates a shadow bufferin the nonlinear journal 300. In state 1234, the global transactionmodule commits transaction T₅. The memory block 180 is in a dirty state,and the shadow buffer in the nonlinear journal 300 is freed. In state1236, the global transaction module requests an order-independentoperation on the respective memory block 180. The respective memoryblock 180 now stores the resultant value, and a delta elementcorresponding to the transaction T₆ is stored in persistent memory 156.

FIG. 12D further illustrates aborting an overwrite of a predecessorblock and its associated delta elements. In 1238, the participant modulereceives a request from the global transaction module to overwrite therespective memory block 180 with the value of “17.” The participantmodule stores the previous value of the memory block by retaining areference to the journal block 190 and the associated delta elements. In1240, the global transaction module aborts transaction T₅. Theparticipant module uses the respective journal block 190 and theassociated delta elements to restore the memory block 180 to itsprevious value “18.” In 1242, the participant module receives a messagefrom the global transaction module to add the value “22” to therespective memory block 180, the participant module adds the value tothe current value in the relevant memory block 180, resulting in thevalue “40.” The participant module also saves a delta elementcorresponding to the order-independent operation corresponding totransaction T₆.

C. Concurrent Transaction Procedure

FIG. 13 illustrates a flowchart of one embodiment implementingconcurrent transactions in a nonlinear journal. In state, 1302, theparticipant module determines whether there is a predecessor journalblock 190 corresponding to the relevant disk block 178. In this example,the relevant disk block 178 is the disk block 178 to which Deltas arebeing applied. In one embodiment, the participant module determineswhether there is a predecessor journal block 190 by traversing thenonlinear journal 300. If there is no predecessor journal block 190, theparticipant module proceeds to state 1307. If there is a predecessorjournal block 190, the participant module proceeds to state 1312. Instate 1307, the participant module performs the relevant operation onthe relevant disk block 178. In this example, the relevant operation isthe order-independent or partially ordered operation for which a Deltais being used. In state 1309, the participant module writes a journalblock 190 that corresponds with the relevant disk block 178, giving itthe data contents of the relevant memory block 180. In this example, therelevant memory block 180 is the memory block 180 corresponding to therelevant disk block 178. If in state 1302, the participant moduledetermines that there is a predecessor journal block 190, theparticipant module proceeds to state 1312. In state 1312, theparticipant module writes a delta element that corresponds to therelevant disk block 178. In state 1314, the participant module 1314performs the relevant operation on the relevant memory block 180.

In the illustrated embodiment, the participant module implements Deltasby looking for a “predecessor block” and performing a delta if one isfound. If there was no predecessor block, the participant module readsthe relevant disk block 178 into a corresponding memory block 180, applythe Delta to the contents of the memory block 180, and write the resultas a full block write. Other embodiments, however, are possible. Inanother embodiment, for example, if there is no predecessor journalblock 190, a local transaction is started. The relevant disk block 178is read for this transaction. A corresponding journal block 190 is thenwritten with the contents of the relevant disk block 178. Thetransaction is committed, and the written journal block 190 becomes thepredecessor journal block 190. Now the Delta may be performed with thepredecessor journal block 190.

FIG. 14 illustrates one embodiment of collapsing delta descriptors innonlinear journal 300. In state 1402, the participant module determineswhether a maximum number of deltas have been written to the nonlinearjournal 300. In one embodiment, the participant module determineswhether a maximum number of deltas has been written for the entirenonlinear journal 300. Additionally and/or alternatively, theparticipant module may determine whether a maximum number of deltas hasbeen written to the nonlinear journal 300 for a particular transaction200. If a maximum number of deltas has been written to the nonlinearjournal 300, the participant module proceeds to state 1404. If a maximumnumber of deltas has not been written to the nonlinear journal 300, theparticipant module proceeds to write the relevant delta to persistentstorage 156. In state 1404, the participant module overwrites therelevant memory block 180 with the combined value of the previouspredecessor journal block 190 and its associated deltas. The participantmodule considers the combined operation as a separate transaction 200,referred to herein as a pseudo-transaction. In state 1406, theparticipant module writes the overwritten relevant memory block 180 topersistent storage 156 as a new predecessor journal block 190. In state1408, the participant module keeps the previous predecessor journalblock 190 in the nonlinear journal 300 as a shadow buffer. In state1410, the participant module determines whether the pseudo-transactionhas committed. If the pseudo-transaction has committed, then theparticipant module proceeds to state 1412. If the pseudo-transaction hasnot committed, the participant module proceeds to state 1414. In state1412, the participant module frees the previous predecessor journalblock 190 and its associated deltas. In state 1414, the participantmodule determines whether the pseudo-transaction aborted. If thepseudo-transaction aborted, the participant module proceeds to state1416. If the pseudo-transaction did not abort, the participant modulereturns to state 1410, and repeats states 1410 and 1414 until thepseudo-transaction commits or aborts. In state 1416, the participantmodule restores the relevant memory block 180 to the combined value ofthe previous predecessor journal block 190 and its deltas.

XI. Other Embodiments

While certain embodiments of the invention have been described, theseembodiments have been presented by way of example only, and are notintended to limit the scope of the present invention. Accordingly, thebreadth and scope of the present invention should be defined inaccordance with the following claims and their equivalents.

1. A concurrent transaction subsystem for a journal as a reliablehigh-speed front end for disk writes, the concurrent transactionsubsystem comprising: a processor; a memory; and a module running on theprocessor, the module configured to: receive a request for a firsttransaction that includes a request to write data in a first location onthe memory; write a first data block corresponding to the firsttransaction to a journal, the journal stored in persistent storage, andthe first data block associated with the first location on the memoryand said first data block in a form suitable for being written to thefirst location on the memory; receive a request for a second transactionthat includes a request to write data in the first location on thememory; and determine whether the second transaction includes an orderindependent operation, wherein an order independent operation is anoperation that will yield the same result regardless of the order ofoperations, and if so, write a first delta element to the journal,wherein the first delta element is representative of a first orderindependent operation and is associated with the first location on thememory, said first delta element being in an intermediate form and notsuitable for directly being written to the first location on the memory.2. The system of claim 1, wherein the first order independent operationincludes at least one of the following: addition, subtraction, integermultiplication, maximum, minimum, bitwise or, set union, and setintersection.
 3. The system of claim 1, wherein the module is furtherconfigured to write a predecessor data structure that is associated withthe first data block, and is further configured to write at least onedelta data structure that is associated with the first delta element. 4.The system of claim 3, wherein the module is further configured totraverse the journal, combing a subset of the predecessor data structureand the at least one delta data structure, and freeing a subset of thepredecessor data structure and the at least one delta data structure. 5.The system of claim 1, wherein the journal system is implemented in afile system.
 6. The system of claim 1, wherein the journal system isimplemented in a distributed system.
 7. The system of claim 1, whereinthe journal system is used to update data that represents at least oneof: accounting, ctime, and parity.
 8. The system of claim 1, wherein themodule is further configured to: receive a request for a thirdtransaction that includes a request to write data in a second locationon the memory; write a second data block corresponding to the thirdwrite transaction to the journal, the second data block associated withthe second location on the memory; receive a request for a fourthtransaction that includes a request to write data in the second locationon the memory; and determine whether the fourth transaction includes apartially ordered operation, wherein a partially ordered operationincludes an operation that is ordered within sub-block boundaries, andif so, write a second delta element to the journal, wherein the seconddelta element is associated with a first partially ordered operation andthe second location on the memory.
 9. The system of claim 8, where thepartially ordered operation includes partial block overwrite.
 10. Thesystem of claim 8, wherein the second delta element is associated withboth a partially ordered and an order independent operation on the sameblock.
 11. A method of implementing a concurrent transaction subsystemfor a journal as a reliable high-speed front end for disk writes, themethod comprising: accessing, by a processor, a request for a firsttransaction that includes a request to write data in a first location ona memory; writing a first data block corresponding to the firsttransaction to a journal, wherein the journal is stored in persistentstorage, and wherein the first data block is associated with the firstlocation on the memory and said first data block is in a form suitablefor being written to the first location on the memory; a request for asecond transaction accessing, by a processor, that includes a request towrite data in the first location on the memory; and determining whetherthe second transaction includes an order independent operation, whereinan order independent operation is an operation that will yield the sameresult regardless of the order of operations, and if so, writing a firstdelta element to the journal stored in persistent storage, wherein thefirst delta element is representative of a first order independentoperation and is associated with the first location on the memory, saidfirst delta element being in an intermediate form and not suitable fordirectly being written to the first location on the memory.
 12. Themethod of claim 11, wherein the first order independent operationincludes at least one of the following: addition, subtraction, integermultiplication, maximum, minimum, bitwise or, set union, and setintersection.
 13. The method of claim 11, wherein the module is furtherconfigured to write a predecessor data structure that is associated withthe first data block, and is further configured to write at least onedelta data structure that is associated with the first delta element.14. The method of claim 13, wherein the module is further configured totraverse the journal, combing a subset of the predecessor data structureand the at least one delta data structure, and freeing a subset of thepredecessor data structure and the at least one delta data structure.15. The method of claim 11, wherein the journal system is implemented ina file system.
 16. The method of claim 11, wherein the journal system isimplemented in a distributed system.
 17. The method of claim 11, whereinthe journal system is used to update data that represents at least oneof: accounting, ctime, and parity.
 18. The method of claim 11, themethod further comprising: accessing, by a processor, a request for athird transaction that includes a request to write data in a secondlocation on the memory; writing a second data block corresponding to thethird transaction to the journal stored in persistent storage, thesecond data block associated with the second location on the memory;accessing, by a processor, a request for a fourth transaction thatincludes a request to write data in the second location on the memory;and determining whether the fourth transaction includes a partiallyordered operation, wherein a partially ordered operation includes anoperation that is ordered within sub-block boundaries, and if so,writing a second delta element to the journal stored in persistentstorage, wherein the second delta element is associated with a firstpartially ordered operation and the second location on the memory. 19.The method of claim 18, where the partially ordered operation includespartial block overwrite.
 20. The method of claim 18, wherein the seconddelta element is associated with both a partially ordered and an orderindependent operation on the same block.